Supporting atomic accumulation with an addressable accumulator

ABSTRACT

Atomically accumulating memory updates in a computer system configured with an accumulator that is memory mapped. The accumulator includes an accumulator memory and an accumulator queue and is configured to communicatively couple to a processor. Included is receiving from the processor, by the accumulator, an accumulation request. The accumulation request includes an accumulation operation identifier and data. Based on determining, by the accumulator, that the accumulator can immediately process the request, immediately processing the request. Processing the request includes atomically updating a value in the accumulator memory, by the accumulator, based on the operation identifier and data of the accumulation request. Based on determining, by the accumulator, that the accumulator is actively processing another accumulation request, queuing, by the accumulator, the accumulation request for later processing. Further included is signaling the processor, by the accumulator, the completion of the accumulation request.

BACKGROUND

This disclosure relates generally to memory and more specifically to amethod, computer program and computer system for atomically accumulatingmemory updates with an addressable accumulator.

The number of central processing unit (CPU) cores on a chip and thenumber of CPU cores connected to a shared memory continues to growsignificantly to support growing workload capacity demand. Theincreasing number of CPUs cooperating to process the same workloads putsa significant burden on software scalability; for example, shared queuesor data-structures protected by traditional semaphores become hot spotsand lead to sub-linear n-way scaling curves. Traditionally this has beencountered by implementing finer-grained locking in software, and withlower latency/higher bandwidth interconnects in hardware. Implementingfine-grained locking to improve software scalability can be verycomplicated and error-prone, and at today's CPU frequencies, thelatencies of hardware interconnects are limited by the physicaldimension of the chips and systems, and by the speed of light.

Implementations of hardware Transactional Memory (HTM, or in thisdiscussion, simply TM) have been introduced, wherein a group ofinstructions—called a transaction—operate in an atomic manner on a datastructure in memory, as viewed by other central processing units (CPUs)and the I/O subsystem (atomic operation is also known as “blockconcurrent” or “serialized” in other literature). The transactionexecutes optimistically without obtaining a lock, but may need to abortand retry the transaction execution if an operation, of the executingtransaction, on a memory location conflicts with another operation onthe same memory location. Previously, software transactional memoryimplementations have been proposed to support software TransactionalMemory (TM). However, hardware TM can provide improved performanceaspects and ease of use over software TM.

U.S. Patent Application Publication US20100235587 A1 titled “StagedSoftware Transactional memory” filed 2009 Mar. 16 and incorporated byreference herein teaches a new form of software transactional memorybased on maps for which data goes through three stages. Updates toshared memory are first redirected to a transaction-private map whichassociates each updated memory location with its transaction-privatevalue. Maps are then added to a shared queue so that multiple versionsof memory can be used concurrently by running transactions. Maps arelater removed from the queue when the updates they refer to have beenapplied to the corresponding memory locations. This design offers a verysimple semantic where starting a transaction takes a stable snapshot ofall transactional objects in memory. It prevents transactions fromaborting or seeing inconsistent data in case of conflict. Performance isinteresting for long running transactions as no synchronization isneeded between a transaction's start and commit, which can themselves belock free.

U.S. Patent Application Publication US20110225376 A1 titled “MemoryManager for a Network Communications Processor Architecture” filed 2010Dec. 9 and incorporated by reference herein teaches describedembodiments provide a memory manager for a network processor having aplurality of processing modules and a shared memory. The memory managerallocates blocks of the shared memory to requesting ones of theplurality of processing modules. A free block list tracks availabilityof memory blocks of the shared memory. A reference counter maintains,for each allocated memory block, a reference count indicating a numberof access requests to the memory block by ones of the plurality ofprocessing modules. The reference count is located with data at theallocated memory block. For subsequent access requests to a given memoryblock concurrent with processing of a prior access request to the memoryblock, a memory access accumulator (i) accumulates an incremental valuecorresponding to the subsequent access requests, (ii) updates thereference count associated with the memory block, and (iii) updates thememory block with the accumulated result.

SUMMARY

According to an embodiment of the disclosure, a computer implementedmethod for atomically accumulating memory updates in a computer systemconfigured with an accumulator that is memory mapped. The accumulatormay include an accumulator memory and an accumulator queue. Theaccumulator may communicatively couple to a processor and may beconfigured to perform a method, said method including: receiving fromthe processor, by the accumulator, an accumulation request directed toan accumulator address. The accumulation request may include anaccumulator operation identifier and data. Based on determining, by theaccumulator, that the accumulator can immediately process the request,immediately processing the request. Processing the request includesatomically updating a value in the accumulator memory, by theaccumulator, based on the operation identifier and data of theaccumulation request. Based on determining, by the accumulator, that theaccumulator is actively processing another accumulation request,queuing, by the accumulator, the accumulation request for laterprocessing. The method may further include signaling the processor, bythe accumulator, a completion of the accumulation request.

In another embodiment of the disclosure, a computer program product foratomically accumulating memory updates in a computer system configuredwith an accumulator that is memory mapped. The computer program productmay include an accumulator having an accumulator memory and anaccumulator queue. The accumulator may communicatively couple to aprocessor and may be configured to perform a method, said methodincluding: receiving from the processor, by the accumulator, anaccumulation request directed to an accumulator address. Theaccumulation request may include an accumulator operation identifier anddata. Based on determining, by the accumulator, that the accumulator canimmediately process the request, immediately processing the request.Processing the request includes atomically updating a value in theaccumulator memory, by the accumulator, based on the operationidentifier and data of the accumulation request. Based on determining,by the accumulator, that the accumulator is actively processing anotheraccumulation request, queuing, by the accumulator, the accumulationrequest for later processing. The method may further include signalingthe processor, by the accumulator, a completion of the accumulationrequest.

In another embodiment of the disclosure, a computer system foratomically accumulating memory updates in a computer system configuredwith an accumulator that is memory mapped. The computer system mayinclude an accumulator having an accumulator memory and an accumulatorqueue. The accumulator may communicatively couple to a processor and maybe configured to perform a method, said method including: receiving fromthe processor, by the accumulator, an accumulation request directed toan accumulator address. The accumulation request may include anaccumulator operation identifier and data. Based on determining, by theaccumulator, that the accumulator can immediately process the request,immediately processing the request. Processing the request includesatomically updating a value in the accumulator memory, by theaccumulator, based on the operation identifier and data of theaccumulation request. Based on determining, by the accumulator, that theaccumulator is actively processing another accumulation request,queuing, by the accumulator, the accumulation request for laterprocessing. The method may further include signaling the processor, bythe accumulator, a completion of the accumulation request.

BRIEF DESCRIPTION OF THE SEVERAL VIEWS OF THE DRAWINGS

One or more aspects of the present disclosed embodiments areparticularly pointed out and distinctly claimed as examples in theclaims at the conclusion of the specification. The foregoing and otherobjects, features, and advantages of the disclosed embodiment areapparent from the following detailed description taken in conjunctionwith the accompanying drawings in which:

FIG. 1 depicts an example multicore transactional memory environment, inaccordance with an illustrative embodiment;

FIG. 2 depicts an example multicore transactional memory environment, inaccordance with an illustrative embodiment;

FIG. 3 depicts example components of an example CPU, in accordance withan illustrative embodiment;

FIG. 4 depicts an exemplary schematic block diagram illustrating amemory mapped addressable accumulator attached to a memory bus and aninterconnect in a multiprocessor environment, within the data processingenvironment of FIG. 10, in accordance with an embodiment of thedisclosure;

FIG. 5 depicts an exemplary schematic block diagram illustrating amemory mapped addressable accumulator and an input accumulator requestwithin the data processing environment of FIG. 10, in accordance with anembodiment of the disclosure;

FIG. 6 is a flowchart illustrating atomic accumulation of memory,illustrated within the data processing environment of FIG. 10, inaccordance with an embodiment of the disclosure;

FIG. 7 depicts an exemplary schematic block diagram illustratingprocessor emulation of a memory mapped addressable accumulator withinthe data processing environment of FIG. 10, in accordance with anembodiment of the disclosure;

FIG. 8 depicts an exemplary schematic block diagram illustratingprocessor emulation of multiple memory mapped addressable accumulatorswithin the data processing environment of FIG. 10, in accordance with anembodiment of the disclosure;

FIG. 9 depicts an exemplary schematic block diagram illustrating oneaccumulator mapping multiple virtual accumulators, within the dataprocessing environment of FIG. 10, in accordance with an embodiment ofthe disclosure;

FIG. 10 is a schematic block diagram which illustrates internal andexternal components of a server computer in accordance with anillustrative embodiment; and

FIGS. 11 and 12 depict an exemplary flow for supporting atomicaccumulation with an addressable accumulator.

DETAILED DESCRIPTION

Historically, a computer system or processor had only a single processor(aka processing unit or central processing unit). The processor includedan instruction processing unit (IPU), a branch unit, a memory controlunit and the like. Such processors were capable of executing a singlethread of a program at a time. Operating systems were developed thatcould time-share a processor by dispatching a program to be executed onthe processor for a period of time, and then dispatching another programto be executed on the processor for another period of time. Astechnology evolved, memory subsystem caches were often added to theprocessor as well as complex dynamic address translation includingtranslation lookaside buffers (TLBs). The IPU itself was often referredto as a processor. As technology continued to evolve, an entireprocessor could be packaged on a single semiconductor chip or die, sucha processor was referred to as a microprocessor. Then processors weredeveloped that incorporated multiple IPUs, such processors were oftenreferred to as multi-processors. Each such processor of amulti-processor computer system (processor) may include individual orshared caches, memory interfaces, system bus, address translationmechanism and the like. Virtual machine and instruction set architecture(ISA) emulators added a layer of software to a processor, that providedthe virtual machine with multiple “virtual processors” (aka processors)by time-slice usage of a single IPU in a single hardware processor. Astechnology further evolved, multi-threaded processors were developed,enabling a single hardware processor having a single multi-thread IPU toprovide a capability of simultaneously executing threads of differentprograms, thus each thread of a multi-threaded processor appeared to theoperating system as a processor. As technology further evolved, it waspossible to put multiple processors (each having an IPU) on a singlesemiconductor chip or die. These processors were referred to processorcores or just cores. Thus the terms such as processor, centralprocessing unit, processing unit, microprocessor, core, processor core,processor thread, and thread, for example, are often usedinterchangeably. Aspects of embodiments herein may be practiced by anyor all processors including those shown supra, without departing fromthe teachings herein. Wherein the term “thread” or “processor thread” isused herein, it is expected that particular advantage of the embodimentmay be had in a processor thread implementation.

Transaction Execution in Intel® Based Embodiments

In “Intel® Architecture Instruction Set Extensions ProgrammingReference” 319433-012A, February 2012, incorporated herein by referencein its entirety, Chapter 8 teaches, in part, that multithreadedapplications may take advantage of increasing numbers of CPU cores toachieve higher performance. However, the writing of multi-threadedapplications requires programmers to understand and take into accountdata sharing among the multiple threads. Access to shared data typicallyrequires synchronization mechanisms. These synchronization mechanismsare used to ensure that multiple threads update shared data byserializing operations that are applied to the shared data, oftenthrough the use of a critical section that is protected by a lock. Sinceserialization limits concurrency, programmers try to limit the overheaddue to synchronization.

Intel® Transactional Synchronization Extensions (Intel® TSX) allow aprocessor to dynamically determine whether threads need to be serializedthrough lock-protected critical sections, and to perform thatserialization only when required. This allows the processor to exposeand exploit concurrency that is hidden in an application because ofdynamically unnecessary synchronization.

With Intel TSX, programmer-specified code regions (also referred to as“transactional regions” or just “transactions”) are executedtransactionally. If the transactional execution completes successfully,then all memory operations performed within the transactional regionwill appear to have occurred instantaneously when viewed from otherprocessors. A processor makes the memory operations of the executedtransaction, performed within the transactional region, visible to otherprocessors only when a successful commit occurs, i.e., when thetransaction successfully completes execution. This process is oftenreferred to as an atomic commit.

Intel TSX provides two software interfaces to specify regions of codefor transactional execution. Hardware Lock Elision (HLE) is a legacycompatible instruction set extension (comprising the XACQUIRE andXRELEASE prefixes) to specify transactional regions. RestrictedTransactional Memory (RTM) is a new instruction set interface(comprising the XBEGIN, XEND, and XABORT instructions) for programmersto define transactional regions in a more flexible manner than thatpossible with HLE. HLE is for programmers who prefer the backwardcompatibility of the conventional mutual exclusion programming model andwould like to run HLE-enabled software on legacy hardware but would alsolike to take advantage of the new lock elision capabilities on hardwarewith HLE support. RTM is for programmers who prefer a flexible interfaceto the transactional execution hardware. In addition, Intel TSX alsoprovides an XTEST instruction. This instruction allows software to querywhether the logical processor is transactionally executing in atransactional region identified by either HLE or RTM.

Since a successful transactional execution ensures an atomic commit, theprocessor executes the code region optimistically without explicitsynchronization. If synchronization was unnecessary for that specificexecution, execution can commit without any cross-thread serialization.If the processor cannot commit atomically, then the optimistic executionfails. When this happens, the processor will roll back the execution, aprocess referred to as a transactional abort. On a transactional abort,the processor will discard all updates performed in the memory regionused by the transaction, restore architectural state to appear as if theoptimistic execution never occurred, and resume executionnon-transactionally.

A processor can perform a transactional abort for numerous reasons. Aprimary reason to abort a transaction is due to conflicting memoryaccesses between the transactionally executing logical processor andanother logical processor. Such conflicting memory accesses may preventa successful transactional execution. Memory addresses read from withina transactional region constitute the read-set of the transactionalregion and addresses written to within the transactional regionconstitute the write-set of the transactional region. Intel TSXmaintains the read- and write-sets at the granularity of a cache line. Aconflicting memory access occurs if another logical processor eitherreads a location that is part of the transactional region's write-set orwrites a location that is a part of either the read- or write-set of thetransactional region. A conflicting access typically means thatserialization is required for this code region. Since Intel TSX detectsdata conflicts at the granularity of a cache line, unrelated datalocations placed in the same cache line will be detected as conflictsthat result in transactional aborts. Transactional aborts may also occurdue to limited transactional resources. For example, the amount of dataaccessed in the region may exceed an implementation-specific capacity.Additionally, some instructions and system events may causetransactional aborts. Frequent transactional aborts result in wastedcycles and increased inefficiency.

Hardware Lock Elision

Hardware Lock Elision (HLE) provides a legacy compatible instruction setinterface for programmers to use transactional execution. HLE providestwo new instruction prefix hints: XACQUIRE and XRELEASE.

With HLE, a programmer adds the XACQUIRE prefix to the front of theinstruction that is used to acquire the lock that is protecting thecritical section. The processor treats the prefix as a hint to elide thewrite associated with the lock acquire operation. Even though the lockacquire has an associated write operation to the lock, the processordoes not add the address of the lock to the transactional region'swrite-set nor does it issue any write requests to the lock. Instead, theaddress of the lock is added to the read-set. The logical processorenters transactional execution. If the lock was available before theXACQUIRE prefixed instruction, then all other processors will continueto see the lock as available afterwards. Since the transactionallyexecuting logical processor neither added the address of the lock to itswrite-set nor performed externally visible write operations to the lock,other logical processors can read the lock without causing a dataconflict. This allows other logical processors to also enter andconcurrently execute the critical section protected by the lock. Theprocessor automatically detects any data conflicts that occur during thetransactional execution and will perform a transactional abort ifnecessary.

Even though the eliding processor did not perform any external writeoperations to the lock, the hardware ensures program order of operationson the lock. If the eliding processor itself reads the value of the lockin the critical section, it will appear as if the processor had acquiredthe lock, i.e. the read will return the non-elided value. This behaviorallows an HLE execution to be functionally equivalent to an executionwithout the HLE prefixes.

An XRELEASE prefix can be added in front of an instruction that is usedto release the lock protecting a critical section. Releasing the lockinvolves a write to the lock. If the instruction is to restore the valueof the lock to the value the lock had prior to the XACQUIRE prefixedlock acquire operation on the same lock, then the processor elides theexternal write request associated with the release of the lock and doesnot add the address of the lock to the write-set. The processor thenattempts to commit the transactional execution.

With HLE, if multiple threads execute critical sections protected by thesame lock but they do not perform any conflicting operations on eachother's data, then the threads can execute concurrently and withoutserialization. Even though the software uses lock acquisition operationson a common lock, the hardware recognizes this, elides the lock, andexecutes the critical sections on the two threads without requiring anycommunication through the lock—if such communication was dynamicallyunnecessary.

If the processor is unable to execute the region transactionally, thenthe processor will execute the region non-transactionally and withoutelision. HLE enabled software has the same forward progress guaranteesas the underlying non-HLE lock-based execution. For successful HLEexecution, the lock and the critical section code must follow certainguidelines. These guidelines only affect performance; and failure tofollow these guidelines will not result in a functional failure.Hardware without HLE support will ignore the XACQUIRE and XRELEASEprefix hints and will not perform any elision since these prefixescorrespond to the REPNE/REPE IA-32 prefixes which are ignored on theinstructions where XACQUIRE and XRELEASE are valid. Importantly, HLE iscompatible with the existing lock-based programming model. Improper useof hints will not cause functional bugs though it may expose latent bugsalready in the code.

Restricted Transactional Memory (RTM) provides a flexible softwareinterface for transactional execution. RTM provides three newinstructions—XBEGIN, XEND, and XABORT—for programmers to start, commit,and abort a transactional execution.

The programmer uses the XBEGIN instruction to specify the start of atransactional code region and the XEND instruction to specify the end ofthe transactional code region. If the RTM region could not besuccessfully executed transactionally, then the XBEGIN instruction takesan operand that provides a relative offset to the fallback instructionaddress.

A processor may abort RTM transactional execution for many reasons. Inmany instances, the hardware automatically detects transactional abortconditions and restarts execution from the fallback instruction addresswith the architectural state corresponding to that present at the startof the XBEGIN instruction and the EAX register updated to describe theabort status.

The XABORT instruction allows programmers to abort the execution of anRTM region explicitly. The XABORT instruction takes an 8-bit immediateargument that is loaded into the EAX register and will thus be availableto software following an RTM abort. RTM instructions do not have anydata memory location associated with them. While the hardware providesno guarantees as to whether an RTM region will ever successfully committransactionally, most transactions that follow the recommendedguidelines are expected to successfully commit transactionally. However,programmers must always provide an alternative code sequence in thefallback path to guarantee forward progress. This may be as simple asacquiring a lock and executing the specified code regionnon-transactionally. Further, a transaction that always aborts on agiven implementation may complete transactionally on a futureimplementation. Therefore, programmers must ensure the code paths forthe transactional region and the alternative code sequence arefunctionally tested.

Detection of HLE Support

A processor supports HLE execution if CPUID.07H.EBX.HLE [bit 4]=1.However, an application can use the HLE prefixes (XACQUIRE and XRELEASE)without checking whether the processor supports HLE. Processors withoutHLE support ignore these prefixes and will execute the code withoutentering transactional execution.

Detection of RTM Support

A processor supports RTM execution if CPUID.07H.EBX.RTM [bit 11]=1. Anapplication must check if the processor supports RTM before it uses theRTM instructions (XBEGIN, XEND, XABORT). These instructions willgenerate a #UD exception when used on a processor that does not supportRTM.

Detection of XTEST Instruction

A processor supports the XTEST instruction if it supports either HLE orRTM. An application must check either of these feature flags beforeusing the XTEST instruction. This instruction will generate a #UDexception when used on a processor that does not support either HLE orRTM.

Querying Transactional Execution Status

The XTEST instruction can be used to determine the transactional statusof a transactional region specified by HLE or RTM. Note, while the HLEprefixes are ignored on processors that do not support HLE, the XTESTinstruction will generate a #UD exception when used on processors thatdo not support either HLE or RTM.

Requirements for HLE Locks

For HLE execution to successfully commit transactionally, the lock mustsatisfy certain properties and access to the lock must follow certainguidelines.

An XRELEASE prefixed instruction must restore the value of the elidedlock to the value it had before the lock acquisition. This allowshardware to safely elide locks by not adding them to the write-set. Thedata size and data address of the lock release (XRELEASE prefixed)instruction must match that of the lock acquire (XACQUIRE prefixed) andthe lock must not cross a cache line boundary.

Software should not write to the elided lock inside a transactional HLEregion with any instruction other than an XRELEASE prefixed instruction,otherwise such a write may cause a transactional abort. In addition,recursive locks (where a thread acquires the same lock multiple timeswithout first releasing the lock) may also cause a transactional abort.Note that software can observe the result of the elided lock acquireinside the critical section. Such a read operation will return the valueof the write to the lock.

The processor automatically detects violations to these guidelines, andsafely transitions to a non-transactional execution without elision.Since Intel TSX detects conflicts at the granularity of a cache line,writes to data collocated on the same cache line as the elided lock maybe detected as data conflicts by other logical processors eliding thesame lock.

Transactional Nesting

Both HLE and RTM support nested transactional regions. However, atransactional abort restores state to the operation that startedtransactional execution: either the outermost XACQUIRE prefixed HLEeligible instruction or the outermost XBEGIN instruction. The processortreats all nested transactions as one transaction.

HLE Nesting and Elision

Programmers can nest HLE regions up to an implementation specific depthof MAX_HLE_NEST_COUNT. Each logical processor tracks the nesting countinternally but this count is not available to software. An XACQUIREprefixed HLE-eligible instruction increments the nesting count, and anXRELEASE prefixed HLE-eligible instruction decrements it. The logicalprocessor enters transactional execution when the nesting count goesfrom zero to one. The logical processor attempts to commit only when thenesting count becomes zero. A transactional abort may occur if thenesting count exceeds MAX_HLE_NEST_COUNT.

In addition to supporting nested HLE regions, the processor can alsoelide multiple nested locks. The processor tracks a lock for elisionbeginning with the XACQUIRE prefixed HLE eligible instruction for thatlock and ending with the XRELEASE prefixed HLE eligible instruction forthat same lock. The processor can, at any one time, track up to aMAX_HLE_ELIDED_LOCKS number of locks. For example, if the implementationsupports a MAX_HLE_ELIDED_LOCKS value of two and if the programmer neststhree HLE identified critical sections (by performing XACQUIRE prefixedHLE eligible instructions on three distinct locks without performing anintervening XRELEASE prefixed HLE eligible instruction on any one of thelocks), then the first two locks will be elided, but the third won't beelided (but will be added to the transaction's writeset). However, theexecution will still continue transactionally. Once an XRELEASE for oneof the two elided locks is encountered, a subsequent lock acquiredthrough the XACQUIRE prefixed HLE eligible instruction will be elided.

The processor attempts to commit the HLE execution when all elidedXACQUIRE and XRELEASE pairs have been matched, the nesting count goes tozero, and the locks have satisfied requirements. If execution cannotcommit atomically, then execution transitions to a non-transactionalexecution without elision as if the first instruction did not have anXACQUIRE prefix.

RTM Nesting

Programmers can nest RTM regions up to an implementation specificMAX_RTM_NEST_COUNT. The logical processor tracks the nesting countinternally but this count is not available to software. An XBEGINinstruction increments the nesting count, and an XEND instructiondecrements the nesting count. The logical processor attempts to commitonly if the nesting count becomes zero. A transactional abort occurs ifthe nesting count exceeds MAX_RTM_NEST_COUNT.

Nesting HLE and RTM

HLE and RTM provide two alternative software interfaces to a commontransactional execution capability. Transactional processing behavior isimplementation specific when HLE and RTM are nested together, e.g., HLEis inside RTM or RTM is inside HLE. However, in all cases, theimplementation will maintain HLE and RTM semantics. An implementationmay choose to ignore HLE hints when used inside RTM regions, and maycause a transactional abort when RTM instructions are used inside HLEregions. In the latter case, the transition from transactional tonon-transactional execution occurs seamlessly since the processor willre-execute the HLE region without actually doing elision, and thenexecute the RTM instructions.

Abort Status Definition

RTM uses the EAX register to communicate abort status to software.Following an RTM abort the EAX register has the following definition.

TABLE 1 RTM Abort Status Definition EAX Register Bit Position Meaning 0Set if abort caused by XABORT instruction 1 If set, the transaction maysucceed on retry, this bit is always clear if bit 0 is set 2 Set ifanother logical processor conflicted with a memory address that was partof the transaction that aborted 3 Set if an internal buffer overflowed 4Set if a debug breakpoint was hit 5 Set if an abort occurred duringexecution of a nested transaction 23:6 Reserved 31-24 XABORT argument(only valid if bit 0 set, otherwise reserved)

The EAX abort status for RTM only provides causes for aborts. It doesnot by itself encode whether an abort or commit occurred for the RTMregion. The value of EAX can be 0 following an RTM abort. For example, aCPUID instruction when used inside an RTM region causes a transactionalabort and may not satisfy the requirements for setting any of the EAXbits. This may result in an EAX value of 0.

RTM Memory Ordering

A successful RTM commit causes all memory operations in the RTM regionto appear to execute atomically. A successfully committed RTM regionconsisting of an XBEGIN followed by an XEND, even with no memoryoperations in the RTM region, has the same ordering semantics as a LOCKprefixed instruction.

The XBEGIN instruction does not have fencing semantics. However, if anRTM execution aborts, then all memory updates from within the RTM regionare discarded and are not made visible to any other logical processor.

RTM-Enabled Debugger Support

By default, any debug exception inside an RTM region will cause atransactional abort and will redirect control flow to the fallbackinstruction address with architectural state recovered and bit 4 in EAXset. However, to allow software debuggers to intercept execution ondebug exceptions, the RTM architecture provides additional capability.

If bit 11 of DR7 and bit 15 of the IA32_DEBUGCTL_MSR are both 1, any RTMabort due to a debug exception (#DB) or breakpoint exception (#BP)causes execution to roll back and restart from the XBEGIN instructioninstead of the fallback address. In this scenario, the EAX register willalso be restored back to the point of the XBEGIN instruction.

Programming Considerations

Typical programmer-identified regions are expected to transactionallyexecute and commit successfully. However, Intel TSX does not provide anysuch guarantee. A transactional execution may abort for many reasons. Totake full advantage of the transactional capabilities, programmersshould follow certain guidelines to increase the probability of theirtransactional execution committing successfully.

This section discusses various events that may cause transactionalaborts. The architecture ensures that updates performed within atransaction that subsequently aborts execution will never becomevisible. Only committed transactional executions initiate an update tothe architectural state. Transactional aborts never cause functionalfailures and only affect performance.

Instruction Based Considerations

Programmers can use any instruction safely inside a transaction (HLE orRTM) and can use transactions at any privilege level. However, someinstructions will always abort the transactional execution and causeexecution to seamlessly and safely transition to a non-transactionalpath.

Intel TSX allows for most common instructions to be used insidetransactions without causing aborts. The following operations inside atransaction do not typically cause an abort:

-   -   Operations on the instruction pointer register, general purpose        registers (GPRs) and the status flags (CF, OF, SF, PF, AF, and        ZF); and    -   Operations on XMM and YMM registers and the MXCSR register.

However, programmers must be careful when intermixing SSE and AVXoperations inside a transactional region. Intermixing SSE instructionsaccessing XMM registers and AVX instructions accessing YMM registers maycause transactions to abort. Programmers may use REP/REPNE prefixedstring operations inside transactions. However, long strings may causeaborts. Further, the use of CLD and STD instructions may cause aborts ifthey change the value of the DF flag. However, if DF is 1, the STDinstruction will not cause an abort. Similarly, if DF is 0, then the CLDinstruction will not cause an abort.

Instructions not enumerated here as causing abort when used inside atransaction will typically not cause a transaction to abort (examplesinclude but are not limited to MFENCE, LFENCE, SFENCE, RDTSC, RDTSCP,etc.).

The following instructions will abort transactional execution on anyimplementation:

XABORT

CPUID

PAUSE

In addition, in some implementations, the following instructions mayalways cause transactional aborts. These instructions are not expectedto be commonly used inside typical transactional regions. However,programmers must not rely on these instructions to force a transactionalabort, since whether they cause transactional aborts is implementationdependent.

-   -   Operations on X87 and MMX architecture state. This includes all        MMX and X87 instructions, including the FXRSTOR and FXSAVE        instructions.    -   Update to non-status portion of EFLAGS: CLI, STI, POPFD, POPFQ,        CLTS.    -   Instructions that update segment registers, debug registers        and/or control registers:    -   MOV to DS/ES/FS/GS/SS, POP DS/ES/FS/GS/SS, LDS, LES, LFS, LGS,        LSS, SWAPGS, WRFSBASE, WRGSBASE, LGDT, SGDT, LIDT, SIDT, LLDT,        SLDT, LTR, STR, Far CALL, Far JMP, Far RET, IRET, MOV to DRx,        MOV to CR0/CR2/CR3/CR4/CR8 and LMSW.    -   Ring transitions: SYSENTER, SYSCALL, SYSEXIT, and SYSRET.    -   TLB and Cacheability control: CLFLUSH, INVD, WBINVD, INVLPG,        INVPCID, and memory instructions with a non-temporal hint        (MOVNTDQA, MOVNTDQ, MOVNTI, MOVNTPD, MOVNTPS, and MOVNTQ).    -   Processor state save: XSAVE, XSAVEOPT, and XRSTOR.    -   Interrupts: INTn, INTO.    -   IO: IN, INS, REP INS, OUT, OUTS, REP OUTS and their variants.    -   VMX: VMPTRLD, VMPTRST, VMCLEAR, VMREAD, VMWRITE, VMCALL,        VMLAUNCH, VMRESUME, VMXOFF, VMXON, INVEPT, and INVVPID.    -   SMX: GETSEC.    -   UD2, RSM, RDMSR, WRMSR, HLT, MONITOR, MWAIT, XSETBV, VZEROUPPER,        MASKMOVQ, and V/MASKMOVDQU.

Runtime Considerations

In addition to the instruction-based considerations, runtime events maycause transactional execution to abort. These may be due to data accesspatterns or micro-architectural implementation features. The followinglist is not a comprehensive discussion of all abort causes.

Any fault or trap in a transaction that must be exposed to software willbe suppressed. Transactional execution will abort and execution willtransition to a non-transactional execution, as if the fault or trap hadnever occurred. If an exception is not masked, then that un-maskedexception will result in a transactional abort and the state will appearas if the exception had never occurred.

Synchronous exception events (#DE, #OF, #NP, #SS, #GP, #BR, #UD, #AC,#XF, #PF, #NM, #TS, #MF, #DB, #BP/INT3) that occur during transactionalexecution may cause an execution not to commit transactionally, andrequire a non-transactional execution. These events are suppressed as ifthey had never occurred. With HLE, since the non-transactional code pathis identical to the transactional code path, these events will typicallyre-appear when the instruction that caused the exception is re-executednon-transactionally, causing the associated synchronous events to bedelivered appropriately in the non-transactional execution. Asynchronousevents (NMI, SMI, INTR, IPI, PMI, etc.) occurring during transactionalexecution may cause the transactional execution to abort and transitionto a non-transactional execution. The asynchronous events will be pendedand handled after the transactional abort is processed.

Transactions only support write-back cacheable memory type operations. Atransaction may always abort if the transaction includes operations onany other memory type. This includes instruction fetches to UC memorytype.

Memory accesses within a transactional region may require the processorto set the Accessed and Dirty flags of the referenced page table entry.The behavior of how the processor handles this is implementationspecific. Some implementations may allow the updates to these flags tobecome externally visible even if the transactional region subsequentlyaborts. Some Intel TSX implementations may choose to abort thetransactional execution if these flags need to be updated. Further, aprocessor's page-table walk may generate accesses to its owntransactionally written but uncommitted state. Some Intel TSXimplementations may choose to abort the execution of a transactionalregion in such situations. Regardless, the architecture ensures that, ifthe transactional region aborts, then the transactionally written statewill not be made architecturally visible through the behavior ofstructures such as TLBs.

Executing self-modifying code transactionally may also causetransactional aborts. Programmers must continue to follow the Intelrecommended guidelines for writing self-modifying and cross-modifyingcode even when employing HLE and RTM. While an implementation of RTM andHLE will typically provide sufficient resources for executing commontransactional regions, implementation constraints and excessive sizesfor transactional regions may cause a transactional execution to abortand transition to a non-transactional execution. The architectureprovides no guarantee of the amount of resources available to dotransactional execution and does not guarantee that a transactionalexecution will ever succeed.

Conflicting requests to a cache line accessed within a transactionalregion may prevent the transaction from executing successfully. Forexample, if logical processor P0 reads line A in a transactional regionand another logical processor P1 writes line A (either inside or outsidea transactional region) then logical processor P0 may abort if logicalprocessor P1's write interferes with processor P0's ability to executetransactionally.

Similarly, if P0 writes line A in a transactional region and P1 reads orwrites line A (either inside or outside a transactional region), then P0may abort if P1's access to line A interferes with P0's ability toexecute transactionally. In addition, other coherence traffic may attimes appear as conflicting requests and may cause aborts. While thesefalse conflicts may happen, they are expected to be uncommon. Theconflict resolution policy to determine whether P0 or P1 aborts in theabove scenarios is implementation specific.

Generic Transaction Execution Embodiments:

According to “ARCHITECTURES FOR TRANSACTIONAL MEMORY”, a dissertationsubmitted to the Department of Computer Science and the Committee onGraduate Studies of Stanford University in partial fulfillment of therequirements for the Degree of Doctor of Philosophy, by Austen McDonald,June 2009, incorporated by reference herein in its entirety,fundamentally, there are three mechanisms needed to implement an atomicand isolated transactional region: versioning, conflict detection, andcontention management.

To make a transactional code region appear atomic, all the modificationsperformed by that transactional code region must be stored and keptisolated from other transactions until commit time. The system does thisby implementing a versioning policy. Two versioning paradigms exist:eager and lazy. An eager versioning system stores newly generatedtransactional values in place and stores previous memory values on theside, in what is called an undo-log. A lazy versioning system stores newvalues temporarily in what is called a write buffer, copying them tomemory only on commit. In either system, the cache is used to optimizestorage of new versions.

To ensure that transactions appear to be performed atomically, conflictsmust be detected and resolved. The two systems, i.e., the eager and lazyversioning systems, detect conflicts by implementing a conflictdetection policy, either optimistic or pessimistic. An optimistic systemexecutes transactions in parallel, checking for conflicts only when atransaction commits. A pessimistic system checks for conflicts at eachload and store. Similar to versioning, conflict detection also uses thecache, marking each line as either part of the read-set, part of thewrite-set, or both. The two systems resolve conflicts by implementing acontention management policy. Many contention management policies exist,some are more appropriate for optimistic conflict detection and some aremore appropriate for pessimistic. Described below are some examplepolicies.

Since each transactional memory (TM) system needs both versioningdetection and conflict detection, these options give rise to fourdistinct TM designs: Eager-Pessimistic (EP), Eager-Optimistic (EO),Lazy-Pessimistic (LP), and Lazy-Optimistic (LO). Table 2 brieflydescribes all four distinct TM designs.

FIGS. 1 and 2 depict an example of a multicore TM environment. FIG. 1shows many TM-enabled CPUs (CPU1 114 a, CPU2 114 b, etc.) on one die100, connected with an interconnect 122, under management of aninterconnect control 120 a, 120 b. Each CPU 114 a, 114 b (also known asa Processor) may have a split cache consisting of an Instruction Cache116 a, 166 b for caching instructions from memory to be executed and aData Cache 118 a, 118 b with TM support for caching data (operands) ofmemory locations to be operated on by CPU 114 a, 114 b (in FIG. 1, eachCPU 114 a, 114 b and its associated caches are referenced as 112 a, 112b). In an implementation, caches of multiple dies 100 are interconnectedto support cache coherency between the caches of the multiple dies 100.In an implementation, a single cache, rather than the split cache isemployed holding both instructions and data. In implementations, the CPUcaches are one level of caching in a hierarchical cache structure. Forexample each die 100 may employ a shared cache 124 to be shared amongstall the CPUs on the die 100. In another implementation, each die mayhave access to a shared cache 124, shared amongst all the processors ofall the dies 100.

FIG. 2 shows the details of an example transactional CPU environment112, having a CPU 114, including additions to support TM. Thetransactional CPU (processor) 114 may include hardware for supportingRegister Checkpoints 126 and special TM Registers 128. The transactionalCPU cache may have the MESI bits 130, Tags 140 and Data 142 of aconventional cache but also, for example, R bits 132 showing a line hasbeen read by the CPU 114 while executing a transaction and W bits 138showing a line has been written-to by the CPU 114 while executing atransaction.

A key detail for programmers in any TM system is how non-transactionalaccesses interact with transactions. By design, transactional accessesare screened from each other using the mechanisms above. However, theinteraction between a regular, non-transactional load with a transactioncontaining a new value for that address must still be considered. Inaddition, the interaction between a non-transactional store with atransaction that has read that address must also be explored. These areissues of the database concept isolation.

A TM system is said to implement strong isolation, sometimes calledstrong atomicity, when every non-transactional load and store acts likean atomic transaction. Therefore, non-transactional loads cannot seeuncommitted data and non-transactional stores cause atomicity violationsin any transactions that have read that address. A system where this isnot the case is said to implement weak isolation, sometimes called weakatomicity.

Strong isolation is often more desirable than weak isolation due to therelative ease of conceptualization and implementation of strongisolation. Additionally, if a programmer has forgotten to surround someshared memory references with transactions, causing bugs, then withstrong isolation, the programmer will often detect that oversight usinga simple debug interface because the programmer will see anon-transactional region causing atomicity violations. Also, programswritten in one model may work differently on another model.

Further, strong isolation is often easier to support in hardware TM thanweak isolation. With strong isolation, since the coherence protocolalready manages load and store communication between processors,transactions can detect non-transactional loads and stores and actappropriately. To implement strong isolation in software TransactionalMemory (TM), non-transactional code must be modified to include read-and write-barriers; potentially crippling performance. Although greateffort has been expended to remove many un-needed barriers, suchtechniques are often complex and performance is typically far lower thanthat of HTMs.

TABLE 2 Transactional Memory Design Space VERSIONING Lazy Eager CONFLICTOptimistic Storing updates in a write Not practical: waiting to updateDETECTION buffer; detecting conflicts at memory until commit time butcommit time. detecting conflicts at access time guarantees wasted workand provides no advantage Pessimistic Storing updates in a Updatingmemory, keeping old writebuffer; detecting values in undo log; detectingconflicts at access time. conflicts at access time.

Table 2 illustrates the fundamental design space of transactional memory(versioning and conflict detection).

Eager-Pessimistic (EP)

This first TM design described below is known as Eager-Pessimistic. AnEP system stores its write-set “in place” (hence the name “eager”) and,to support rollback, stores the old values of overwritten lines in an“undo log”. Processors use the W 138 and R 132 cache bits to track readand write-sets and detect conflicts when receiving snooped loadrequests. Perhaps the most notable examples of EP systems in knownliterature are Log TM and UTM.

Beginning a transaction in an EP system is much like beginning atransaction in other systems: tm_begin( ) takes a register checkpoint,and initializes any status registers. An EP system also requiresinitializing the undo log, the details of which are dependent on the logformat, but often involve initializing a log base pointer to a region ofpre-allocated, thread-private memory, and clearing a log boundsregister.

Versioning: In EP, due to the way eager versioning is designed tofunction, the MESI 130 state transitions (cache line indicatorscorresponding to Modified, Exclusive, Shared, and Invalid code states)are left mostly unchanged. Outside of a transaction, the MESI 130 statetransitions are left completely unchanged. When reading a line inside atransaction, the standard coherence transitions apply (S (Shared)→S, I(Invalid)→S, or I→E (Exclusive)), issuing a load miss as needed, but theR 132 bit is also set. Likewise, writing a line applies the standardtransitions (S→M, E→I, I→M), issuing a miss as needed, but also sets theW 138 (Written) bit. The first time a line is written, the old versionof the entire line is loaded then written to the undo log to preserve itin case the current transaction aborts. The newly written data is thenstored “in-place,” over the old data.

Conflict Detection: Pessimistic conflict detection uses coherencemessages exchanged on misses, or upgrades, to look for conflicts betweentransactions. When a read miss occurs within a transaction, otherprocessors receive a load request; but they ignore the request if theydo not have the needed line. If the other processors have the neededline non-speculatively or have the line R 132 (Read), they downgradethat line to S, and in certain cases issue a cache-to-cache transfer ifthey have the line in MESI's 130 M or E state. However, if the cache hasthe line W 138, then a conflict is detected between the two transactionsand additional action(s) must be taken.

Similarly, when a transaction seeks to upgrade a line from shared tomodified (on a first write), the transaction issues an exclusive loadrequest, which is also used to detect conflicts. If a receiving cachehas the line non-speculatively, then the line is invalidated, and incertain cases a cache-to-cache transfer (M or E states) is issued. But,if the line is R 132 or W 138, a conflict is detected.

Validation: Because conflict detection is performed on every load, atransaction always has exclusive access to its own write-set. Therefore,validation does not require any additional work.

Commit: Since eager versioning stores the new version of data items inplace, the commit process simply clears the W 138 and R 132 bits anddiscards the undo log.

Abort: When a transaction rolls back, the original version of each cacheline in the undo log must be restored, a process called “unrolling” or“applying” the log. This is done during tm_discard( ) and must be atomicwith regard to other transactions. Specifically, the write-set muststill be used to detect conflicts: this transaction has the only correctversion of lines in its undo log, and requesting transactions must waitfor the correct version to be restored from that log. Such a log can beapplied using a hardware state machine or software abort handler.

Eager-Pessimistic has the characteristics of: Commit is simple and sinceit is in-place, very fast. Similarly, validation is a no-op. Pessimisticconflict detection detects conflicts early, thereby reducing the numberof “doomed” transactions. For example, if two transactions are involvedin a Write-After-Read dependency, then that dependency is detectedimmediately in pessimistic conflict detection. However, in optimisticconflict detection such conflicts are not detected until the writercommits.

Eager-Pessimistic also has the characteristics of: As described above,the first time a cache line is written, the old value must be written tothe log, incurring extra cache accesses. Aborts are expensive as theyrequire undoing the log. For each cache line in the log, a load must beissued, perhaps going as far as main memory before continuing to thenext line. Pessimistic conflict detection also prevents certainserializable schedules from existing.

Additionally, because conflicts are handled as they occur, there is apotential for livelock and careful contention management mechanisms mustbe employed to guarantee forward progress.

Lazy-Optimistic (LO)

Another popular TM design is Lazy-Optimistic (LO), which stores itswrite-set in a “write buffer” or “redo log” and detects conflicts atcommit time (still using the R 132 and W 138 bits).

Versioning: Just as in the EP system, the MESI protocol of the LO designis enforced outside of the transactions. Once inside a transaction,reading a line incurs the standard MESI transitions but also sets the R132 bit. Likewise, writing a line sets the W 138 bit of the line, buthandling the MESI transitions of the LO design is different from that ofthe EP design. First, with lazy versioning, the new versions of writtendata are stored in the cache hierarchy until commit while othertransactions have access to old versions available in memory or othercaches. To make available the old versions, dirty lines (M lines) mustbe evicted when first written by a transaction. Second, no upgrademisses are needed because of the optimistic conflict detection feature:if a transaction has a line in the S state, it can simply write to itand upgrade that line to an M state without communicating the changeswith other transactions because conflict detection is done at committime.

Conflict Detection and Validation: To validate a transaction and detectconflicts, LO communicates the addresses of speculatively modified linesto other transactions only when it is preparing to commit. Onvalidation, the processor sends one, potentially large, network packetcontaining all the addresses in the write-set. Data is not sent, butleft in the cache of the committer and marked dirty (M). To build thispacket without searching the cache for lines marked W, a simple bitvector is used, called a “store buffer,” with one bit per cache line totrack these speculatively modified lines. Other transactions use thisaddress packet to detect conflicts: if an address is found in the cacheand the R 132 and/or W 138 bits are set, then a conflict is initiated.If the line is found but neither R 132 nor W 138 is set, then the lineis simply invalidated, which is similar to processing an exclusive load.

To support transaction atomicity, these address packets must be handledatomically, i.e., no two address packets may exist at once with the sameaddresses. In an LO system, this can be achieved by simply acquiring aglobal commit token before sending the address packet. However, atwo-phase commit scheme could be employed by first sending out theaddress packet, collecting responses, enforcing an ordering protocol(perhaps oldest transaction first), and committing once all responsesare satisfactory.

Commit: Once validation has occurred, commit needs no special treatment:simply clear W 138 and R 132 bits and the store buffer. Thetransaction's writes are already marked dirty in the cache and othercaches' copies of these lines have been invalidated via the addresspacket. Other processors can then access the committed data through theregular coherence protocol.

Abort: Rollback is equally easy: because the write-set is containedwithin the local caches, these lines can be invalidated, then clear W138 and R 132 bits and the store buffer. The store buffer allows W linesto be found to invalidate without the need to search the cache.

Lazy-Optimistic has the characteristics of: Aborts are very fast,requiring no additional loads or stores and making only local changes.More serializable schedules can exist than found in EP, which allows anLO system to more aggressively speculate that transactions areindependent, which can yield higher performance. Finally, the latedetection of conflicts can increase the likelihood of forward progress.

Lazy-Optimistic also has the characteristics of: Validation takes globalcommunication time proportional to size of write set. Doomedtransactions can waste work since conflicts are detected only at committime.

Lazy-Pessimistic (LP)

Lazy-Pessimistic (LP) represents a third TM design option, sittingsomewhere between EP and LO: storing newly written lines in a writebuffer but detecting conflicts on a per access basis.

Versioning: Versioning is similar but not identical to that of LO:reading a line sets its R bit 132, writing a line sets its W bit 138,and a store buffer is used to track W lines in the cache. Also, dirty(M) lines must be evicted when first written by a transaction, just asin LO. However, since conflict detection is pessimistic, load exclusivesmust be performed when upgrading a transactional line from I, S→M, whichis unlike LO.

Conflict Detection: LP's conflict detection operates the same as EP's:using coherence messages to look for conflicts between transactions.

Validation: Like in EP, pessimistic conflict detection ensures that atany point, a running transaction has no conflicts with any other runningtransaction, so validation is a no-op.

Commit: Commit needs no special treatment: simply clear W 138 and R 132bits and the store buffer, like in LO.

Abort: Rollback is also like that of LO: simply invalidate the write-setusing the store buffer and clear the W and R bits and the store buffer.

Eager-Optimistic (EO)

The LP has the characteristics of: Like LO, aborts are very fast. LikeEP, the use of pessimistic conflict detection reduces the number of“doomed” transactions. Like EP, some serializable schedules are notallowed and conflict detection must be performed on each cache miss.

The final combination of versioning and conflict detection isEager-Optimistic (EO). EO may be a less than optimal choice for HTMsystems: since new transactional versions are written in-place, othertransactions have no choice but to notice conflicts as they occur (i.e.,as cache misses occur). But since EO waits until commit time to detectconflicts, those transactions become “zombies,” continuing to execute,wasting resources, yet are “doomed” to abort.

EO has proven to be useful in STMs and is implemented by Bartok-STM andMcRT. A lazy versioning STM needs to check its write buffer on each readto ensure that it is reading the most recent value. Since the writebuffer is not a hardware structure, this is expensive, hence thepreference for write-in-place eager versioning. Additionally, sincechecking for conflicts is also expensive in an STM, optimistic conflictdetection offers the advantage of performing this operation in bulk.

Contention Management

How a transaction rolls back once the system has decided to abort thattransaction has been described above, but, since a conflict involves twotransactions, the topics of which transaction should abort, how thatabort should be initiated, and when should the aborted transaction beretried need to be explored. These are topics that are addressed byContention Management (CM), a key component of transactional memory.Described below are policies regarding how the systems initiate abortsand the various established methods of managing which transactionsshould abort in a conflict.

Contention Management Policies

A Contention Management (CM) Policy is a mechanism that determines whichtransaction involved in a conflict should abort and when the abortedtransaction should be retried. For example, it is often the case thatretrying an aborted transaction immediately does not lead to the bestperformance. Conversely, employing a back-off mechanism, which delaysthe retrying of an aborted transaction, can yield better performance.STMs first grappled with finding the best contention management policiesand many of the policies outlined below were originally developed forSTMs.

CM Policies draw on a number of measures to make decisions, includingages of the transactions, size of read- and write-sets, the number ofprevious aborts, etc. The combinations of measures to make suchdecisions are endless, but certain combinations are described below,roughly in order of increasing complexity.

To establish some nomenclature, first note that in a conflict there aretwo sides: the attacker and the defender. The attacker is thetransaction requesting access to a shared memory location. Inpessimistic conflict detection, the attacker is the transaction issuingthe load or load exclusive. In optimistic, the attacker is thetransaction attempting to validate. The defender in both cases is thetransaction receiving the attacker's request.

An Aggressive CM Policy immediately and always retries either theattacker or the defender. In LO, Aggressive means that the attackeralways wins, and so Aggressive is sometimes called committer wins. Sucha policy was used for the earliest LO systems. In the case of EP,Aggressive can be either defender wins or attacker wins.

Restarting a conflicting transaction that will immediately experienceanother conflict is bound to waste work—namely interconnect bandwidthrefilling cache misses. A Polite CM Policy employs exponential backoff(but linear could also be used) before restarting conflicts. To preventstarvation, a situation where a process does not have resourcesallocated to it by the scheduler, the exponential backoff greatlyincreases the odds of transaction success after some n retries.

Another approach to conflict resolution is to randomly abort theattacker or defender (a policy called Randomized). Such a policy may becombined with a randomized backoff scheme to avoid unneeded contention.

However, making random choices, when selecting a transaction to abort,can result in aborting transactions that have completed “a lot of work”,which can waste resources. To avoid such waste, the amount of workcompleted on the transaction can be taken into account when determiningwhich transaction to abort. One measure of work could be a transaction'sage. Other methods include Oldest, Bulk TM, Size Matters, Karma, andPolka. Oldest is a simple timestamp method that aborts the youngertransaction in a conflict. Bulk TM uses this scheme. Size Matters islike Oldest but instead of transaction age, the number of read/writtenwords is used as the priority, reverting to Oldest after a fixed numberof aborts. Karma is similar, using the size of the write-set aspriority. Rollback then proceeds after backing off a fixed amount oftime. Aborted transactions keep their priorities after being aborted(hence the name Karma). Polka works like Karma but instead of backingoff a predefined amount of time, it backs off exponentially more eachtime.

Since aborting wastes work, it is logical to argue that stalling anattacker until the defender has finished their transaction would lead tobetter performance. Unfortunately, such a simple scheme easily leads todeadlock.

Deadlock avoidance techniques can be used to solve this problem. Greedyuses two rules to avoid deadlock. The first rule is, if a firsttransaction, T1, has lower priority than a second transaction, T0, or ifT1 is waiting for another transaction, then T1 aborts when conflictingwith T0. The second rule is, if T1 has higher priority than T0 and isnot waiting, then T0 waits until T1 commits, aborts, or starts waiting(in which case the first rule is applied). Greedy provides someguarantees about time bounds for executing a set of transactions. One EPdesign (Log TM) uses a CM policy similar to Greedy to achieve stallingwith conservative deadlock avoidance.

Example MESI coherency rules provide for four possible states in which acache line of a multiprocessor cache system may reside, M, E, S, and I,defined as follows:

Modified (M): The cache line is present only in the current cache, andis dirty; it has been modified from the value in main memory. The cacheis required to write the data back to main memory at some time in thefuture, before permitting any other read of the (no longer valid) mainmemory state. The write-back changes the line to the Exclusive state.

Exclusive (E): The cache line is present only in the current cache, butis clean; it matches main memory. It may be changed to the Shared stateat any time, in response to a read request. Alternatively, it may bechanged to the Modified state when writing to it.

Shared (S): Indicates that this cache line may be stored in other cachesof the machine and is “clean”; it matches the main memory. The line maybe discarded (changed to the Invalid state) at any time.

Invalid (I): Indicates that this cache line is invalid (unused).

TM coherency status indicators (R 132, W 138) may be provided for eachcache line, in addition to, or encoded in the MESI coherency bits. An R132 indicator indicates the current transaction has read from the dataof the cache line, and a W 138 indicator indicates the currenttransaction has written to the data of the cache line.

In another aspect of TM design, a system is designed using transactionalstore buffers. U.S. Pat. No. 6,349,361 titled “Methods and Apparatus forReordering and Renaming Memory References in a Multiprocessor ComputerSystem,” filed Mar. 31, 2000 and incorporated by reference herein in itsentirety, teaches a method for reordering and renaming memory referencesin a multiprocessor computer system having at least a first and a secondprocessor. The first processor has a first private cache and a firstbuffer, and the second processor has a second private cache and a secondbuffer. The method includes the steps of, for each of a plurality ofgated store requests received by the first processor to store a datum,exclusively acquiring a cache line that contains the datum by the firstprivate cache, and storing the datum in the first buffer. Upon the firstbuffer receiving a load request from the first processor to load aparticular datum, the particular datum is provided to the firstprocessor from among the data stored in the first buffer based on anin-order sequence of load and store operations. Upon the first cachereceiving a load request from the second cache for a given datum, anerror condition is indicated and a current state of at least one of theprocessors is reset to an earlier state when the load request for thegiven datum corresponds to the data stored in the first buffer.

The main implementation components of one such transactional memoryfacility are a transaction-backup register file for holdingpre-transaction GR (general register) content, a cache directory totrack the cache lines accessed during the transaction, a store cache tobuffer stores until the transaction ends, and firmware routines toperform various complex functions. In this section a detailedimplementation is described.

IBM zEnterprise EC12 Enterprise Server Embodiment

The IBM zEnterprise EC12 enterprise server introduces transactionalexecution (TX) in transactional memory, and is described in part in apaper, “Transactional Memory Architecture and Implementation for IBMSystem z” of Proceedings Pages 25-36 presented at MICRO-45, 1-5 Dec.2012, Vancouver, British Columbia, Canada, available from IEEE ComputerSociety Conference Publishing Services (CPS), which is incorporated byreference herein in its entirety.

Table 3 shows an example transaction. Transactions started with TBEGINare not assured to ever successfully complete with TEND, since they canexperience an aborting condition at every attempted execution, e.g., dueto repeating conflicts with other CPUs. This requires that the programsupport a fallback path to perform the same operationnon-transactionally, e.g., by using traditional locking schemes. Thisputs significant burden on the programming and software verificationteams, especially where the fallback path is not automatically generatedby a reliable compiler.

TABLE 3 Example Transaction Code LHI R0,0 *initialize retry count = 0loop TBEGIN *begin transaction JNZ abort *go to abort code if CC1 = 0 LTR1, lock *load and test the fallback lock JNZ lckbzy *branch if lockbusy . . . perform operation . . . TEND *end transaction . . . . . . . .. . . . lckbzy TABORT *abort if lock busy; this *resumes after TBEGINabort JO fallback *no retry if CC = 3 AHI R0, 1 *increment retry countCIJNL R0,6, fallback *give up after 6 attempts PPA R0, TX *random delaybased on retry count . . . potentially wait for lock to become free . .. J loop *jump back to retry fallback OBTAIN lock *using Compare&Swap .. . perform operation . . . RELEASE lock . . . . . . . . . . . .

The requirement of providing a fallback path for aborted TransactionExecution (TX) transactions can be onerous. Many transactions operatingon shared data structures are expected to be short, touch only a fewdistinct memory locations, and use simple instructions only. For thosetransactions, the IBM zEnterprise EC12 introduces the concept ofconstrained transactions; under normal conditions, the CPU 114 (FIG. 2)assures that constrained transactions eventually end successfully,albeit without giving a strict limit on the number of necessary retries.A constrained transaction starts with a TBEGINC instruction and endswith a regular TEND. Implementing a task as a constrained ornon-constrained transaction typically results in very comparableperformance, but constrained transactions simplify software developmentby removing the need for a fallback path. IBM's Transactional Executionarchitecture is further described in z/Architecture, Principles ofOperation, Tenth Edition, SA22-7832-09 published September 2012 fromIBM, incorporated by reference herein in its entirety.

A constrained transaction starts with the TBEGINC instruction. Atransaction initiated with TBEGINC must follow a list of programmingconstraints; otherwise the program takes a non-filterableconstraint-violation interruption. Exemplary constraints may include,but not be limited to: the transaction can execute a maximum of 32instructions, all instruction text must be within 256 consecutive bytesof memory; the transaction contains only forward-pointing relativebranches (i.e., no loops or subroutine calls); the transaction canaccess a maximum of 4 aligned octowords (an octoword is 32 bytes) ofmemory; and restriction of the instruction-set to exclude complexinstructions like decimal or floating-point operations. The constraintsare chosen such that many common operations like doubly linkedlist-insert/delete operations can be performed, including the verypowerful concept of atomic compare-and-swap targeting up to 4 alignedoctowords. At the same time, the constraints were chosen conservativelysuch that future CPU implementations can assure transaction successwithout needing to adjust the constraints, since that would otherwiselead to software incompatibility.

TBEGINC mostly behaves like XBEGIN in TSX or TBEGIN on IBM's zEC12servers, except that the floating-point register (FPR) control and theprogram interruption filtering fields do not exist and the controls areconsidered to be zero. On a transaction abort, the instruction addressis set back directly to the TBEGINC instead of to the instruction after,reflecting the immediate retry and absence of an abort path forconstrained transactions.

Nested transactions are not allowed within constrained transactions, butif a TBEGINC occurs within a non-constrained transaction it is treatedas opening a new non-constrained nesting level just like TBEGIN would.This can occur, e.g., if a non-constrained transaction calls asubroutine that uses a constrained transaction internally.

Since interruption filtering is implicitly off, all exceptions during aconstrained transaction lead to an interruption into the operatingsystem (OS). Eventual successful finishing of the transaction relies onthe capability of the OS to page-in the at most 4 pages touched by anyconstrained transaction. The OS must also ensure time-slices long enoughto allow the transaction to complete.

TABLE 4 Transaction Code Example TBEGINC *begin constrained transaction. . . perform operation . . . TEND *end transaction

Table 4 shows the constrained-transactional implementation of the codein Table 3, assuming that the constrained transactions do not interactwith other locking-based code. No lock testing is shown therefore, butcould be added if constrained transactions and lock-based code weremixed.

When failure occurs repeatedly, software emulation is performed usingmillicode as part of system firmware. Advantageously, constrainedtransactions have desirable properties because of the burden removedfrom programmers.

With reference to FIG. 3, the IBM zEnterprise EC12 processor introducedthe transactional execution facility. The processor can decode 3instructions per clock cycle; simple instructions are dispatched assingle micro-ops, and more complex instructions are cracked intomultiple micro-ops. The micro-ops (Uops 232 b) are written into aunified issue queue 216, from where they can be issued out-of-order. Upto two fixed-point, one floating-point, two load/store, and two branchinstructions can execute every cycle. A Global Completion Table (GCT)232 holds every micro-op 232 b and a transaction nesting depth (TND) 232a. The GCT 232 is written in-order at decode time, tracks the executionstatus of each micro-op 232 b, and completes instructions when allmicro-ops 232 b of the oldest instruction group have successfullyexecuted.

The level 1 (L1) data cache 240 is a 96 KB (kilo-byte) 6-way associativecache with 256 byte cache-lines and 4 cycle use latency, coupled to aprivate 1 MB (mega-byte) 8-way associative 2nd-level (L2) data cache 268with 7 cycles use-latency penalty for L1 240 misses. The L1 240 cache isthe cache closest to a processor and Ln cache is a cache at the nthlevel of caching. Both L1 240 and L2 268 caches are store-through. Sixcores on each central processor (CP) chip share a 48 MB 3rd-levelstore-in cache, and six CP chips are connected to an off-chip 384 MB4th-level cache, packaged together on a glass ceramic multi-chip module(MCM). Up to 4 multi-chip modules (MCMs) can be connected to a coherentsymmetric multi-processor (SMP) system with up to 144 cores (not allcores are available to run customer workload).

Coherency is managed with a variant of the MESI protocol. Cache-linescan be owned read-only (shared) or exclusive; the L1 240 and L2 268 arestore-through and thus do not contain dirty lines. The L3 272 and L4caches (not shown) are store-in and track dirty states. Each cache isinclusive of all its connected lower level caches.

Coherency requests are called “cross interrogates” (XI) and are senthierarchically from higher level to lower-level caches, and between theL4s. When one core misses the L1 240 and L2 268 and requests the cacheline from its local L3 272, the L3 272 checks whether it owns the line,and if necessary sends an XI to the currently owning L2 268/L1 240 underthat L3 272 to ensure coherency, before it returns the cache line to therequestor. If the request also misses the L3 272, the L3 272 sends arequest to the L4 (not shown), which enforces coherency by sending XIsto all necessary L3s under that L4, and to the neighboring L4s. Then theL4 responds to the requesting L3 which forwards the response to the L2268/L1 240.

Note that due to the inclusivity rule of the cache hierarchy, sometimescache lines are XI'ed from lower-level caches due to evictions onhigher-level caches caused by associativity overflows from requests toother cache lines. These XIs can be called “LRU XIs”, where LRU standsfor least recently used.

Making reference to yet another type of XI requests, Demote-XIstransition cache-ownership from exclusive into read-only state, andExclusive-XIs transition cache ownership from exclusive into invalidstate. Demote-XIs and Exclusive-XIs need a response back to the XIsender. The target cache can “accept” the XI, or send a “reject”response if it first needs to evict dirty data before accepting the XI.The L1 240/L2 268 caches are store through, but may reject demote-XIsand exclusive XIs if they have stores in their store queues that need tobe sent to L3 before downgrading the exclusive state. A rejected XI willbe repeated by the sender. Read-only-XIs are sent to caches that own theline read-only; no response is needed for such XIs since they cannot berejected. The details of the SMP protocol are similar to those describedfor the IBM z10 by P. Mak, C. Walters, and G. Strait, in “IBM System z10processor cache subsystem microarchitecture”, IBM Journal of Researchand Development, Vol 53:1, 2009, which is incorporated by referenceherein in its entirety.

Transactional Instruction Execution

FIG. 3 depicts example components of an example transactional executionenvironment, including a CPU and caches/components with which itinteracts (such as those depicted in FIGS. 1 and 2). The instructiondecode unit 208 (IDU) keeps track of the current transaction nestingdepth 212 (TND). When the IDU 208 receives a TBEGIN instruction, thenesting depth 212 is incremented, and conversely decremented on TENDinstructions. The nesting depth 212 is written into the GCT 232 forevery dispatched instruction. When a TBEGIN or TEND is decoded on aspeculative path that later gets flushed, the IDU's 208 nesting depth212 is refreshed from the youngest GCT 232 entry that is not flushed.The transactional state is also written into the issue queue 216 forconsumption by the execution units, mostly by the Load/Store Unit (LSU)280, which also has an effective address calculator 236 is included inthe LSU 280. The TBEGIN instruction may specify a transaction diagnosticblock (TDB) for recording status information, should the transactionabort before reaching a TEND instruction.

Similar to the nesting depth, the IDU 208/GCT 232 collaboratively trackthe access register/floating-point register (AR/FPR) modification masksthrough the transaction nest; the IDU 208 can place an abort requestinto the GCT 232 when an AR/FPR-modifying instruction is decoded and themodification mask blocks that. When the instruction becomesnext-to-complete, completion is blocked and the transaction aborts.Other restricted instructions are handled similarly, including TBEGIN ifdecoded while in a constrained transaction, or exceeding the maximumnesting depth.

An outermost TBEGIN is cracked into multiple micro-ops depending on theGR-Save-Mask; each micro-op 232 b (including, for example uop 0, uop 1,and uop2) will be executed by one of the two fixed point units (FXUs)220 to save a pair of GRs 228 into a special transaction-backup registerfile 224, that is used to later restore the GR 228 content in case of atransaction abort. Also the TBEGIN spawns micro-ops 232 b to perform anaccessibility test for the TDB if one is specified; the address is savedin a special purpose register for later usage in the abort case. At thedecoding of an outermost TBEGIN, the instruction address and theinstruction text of the TBEGIN are also saved in special purposeregisters for a potential abort processing later on.

TEND and NTSTG are single micro-op 232 b instructions; NTSTG(non-transactional store) is handled like a normal store except that itis marked as non-transactional in the issue queue 216 so that the LSU280 can treat it appropriately. TEND is a no-op at execution time, theending of the transaction is performed when TEND completes.

As mentioned, instructions that are within a transaction are marked assuch in the issue queue 216, but otherwise execute mostly unchanged; theLSU 280 performs isolation tracking as described in the next section.

Since decoding is in-order, and since the IDU 208 keeps track of thecurrent transactional state and writes it into the issue queue 216 alongwith every instruction from the transaction, execution of TBEGIN, TEND,and instructions before, within, and after the transaction can beperformed out-of order. It is even possible (though unlikely) that TENDis executed first, then the entire transaction, and lastly the TBEGINexecutes. Program order is restored through the GCT 232 at completiontime. The length of transactions is not limited by the size of the GCT232, since general purpose registers (GRs) 228 can be restored from thebackup register file 224.

During execution, the program event recording (PER) events are filteredbased on the Event Suppression Control, and a PER TEND event is detectedif enabled. Similarly, while in transactional mode, a pseudo-randomgenerator may be causing the random aborts as enabled by the TransactionDiagnostics Control.

Tracking for Transactional Isolation

The Load/Store Unit 280 tracks cache lines that were accessed duringtransactional execution, and triggers an abort if an XI from another CPU(or an LRU-XI) conflicts with the footprint. If the conflicting XI is anexclusive or demote XI, the LSU 280 rejects the XI back to the L3 272 inthe hope of finishing the transaction before the L3 272 repeats the XI.This “stiff-arming” is very efficient in highly contended transactions.In order to prevent hangs when two CPUs stiff-arm each other, aXI-reject counter is implemented, which triggers a transaction abortwhen a threshold is met.

The L1 cache directory 240 is traditionally implemented with staticrandom access memories (SRAMs). For the transactional memoryimplementation, the valid bits 244 (64 rows×6 ways) of the directoryhave been moved into normal logic latches, and are supplemented with twomore bits per cache line: the TX-read 248 and TX-dirty 252 bits.

The TX-read 248 bits are reset when a new outermost TBEGIN is decoded(which is interlocked against a prior still pending transaction). TheTX-read 248 bit is set at execution time by every load instruction thatis marked “transactional” in the issue queue. Note that this can lead toover-marking if speculative loads are executed, for example on amispredicted branch path. The alternative of setting the TX-read 248 bitat load completion time was too expensive for silicon area, sincemultiple loads can complete at the same time, requiring many read-portson the load-queue.

Stores execute the same way as in non-transactional mode, but atransaction mark is placed in the store queue (STQ) 260 entry of thestore instruction. At write-back time, when the data from the STQ 260 iswritten into the L1 240, the TX-dirty bit 252 in the L1-directory 256 isset for the written cache line. Store write-back into the L1 240 occursonly after the store instruction has completed, and at most one store iswritten back per cycle. Before completion and write-back, loads canaccess the data from the STQ 260 by means of store-forwarding; afterwrite-back, the CPU 114 (FIG. 2) can access the speculatively updateddata in the L1 240. If the transaction ends successfully, the TX-dirtybits 252 of all cache-lines are cleared, and also the TX-marks of notyet written stores are cleared in the STQ 260, effectively turning thepending stores into normal stores.

On a transaction abort, all pending transactional stores are invalidatedfrom the STQ 260, even those already completed. All cache lines thatwere modified by the transaction in the L1 240, that is, have theTX-dirty bit 252 on, have their valid bits turned off, effectivelyremoving them from the L1 240 cache instantaneously.

The architecture requires that before completing a new instruction, theisolation of the transaction read- and write-set is maintained. Thisisolation is ensured by stalling instruction completion at appropriatetimes when XIs are pending; speculative out-of order execution isallowed, optimistically assuming that the pending XIs are to differentaddresses and not actually cause a transaction conflict. This designfits very naturally with the XI-vs-completion interlocks that areimplemented on prior systems to ensure the strong memory ordering thatthe architecture requires.

When the L1 240 receives an XI, L1 240 accesses the directory to checkvalidity of the XI'ed address in the L1 240, and if the TX-read bit 248is active on the XI'ed line and the XI is not rejected, the LSU 280triggers an abort. When a cache line with active TX-read bit 248 isLRU'ed from the L1 240, a special LRU-extension vector remembers foreach of the 64 rows of the L1 240 that a TX-read line existed on thatrow. Since no precise address tracking exists for the LRU extensions,any non-rejected XI that hits a valid extension row the LSU 280 triggersan abort. Providing the LRU-extension effectively increases the readfootprint capability from the L1-size to the L2-size and associativity,provided no conflicts with other CPUs 114 (FIGS. 1 and 2) against thenon-precise LRU-extension tracking causes aborts.

The store footprint is limited by the store cache size (the store cacheis discussed in more detail below) and thus implicitly by the L2 268size and associativity. No LRU-extension action needs to be performedwhen a TX-dirty 252 cache line is LRU'ed from the L1 240.

Store Cache

In prior systems, since the L1 240 and L2 268 are store-through caches,every store instruction causes an L3 272 store access; with now 6 coresper L3 272 and further improved performance of each core, the store ratefor the L3 272 (and to a lesser extent for the L2 268) becomesproblematic for certain workloads. In order to avoid store queuingdelays, a gathering store cache 264 had to be added, that combinesstores to neighboring addresses before sending them to the L3 272.

For transactional memory performance, it is acceptable to invalidateevery TX-dirty 252 cache line from the L1 240 on transaction aborts,because the L2 268 cache is very close (7 cycles L1 240 miss penalty) tobring back the clean lines. However, it would be unacceptable forperformance (and silicon area for tracking) to have transactional storeswrite the L2 268 before the transaction ends and then invalidate alldirty L2 268 cache lines on abort (or even worse on the shared L3 272).

The two problems of store bandwidth and transactional memory storehandling can both be addressed with the gathering store cache 264. Thecache 264 is a circular queue of 64 entries, each entry holding 128bytes of data with byte-precise valid bits. In non-transactionaloperation, when a store is received from the LSU 280, the store cache264 checks whether an entry exists for the same address, and if sogathers the new store into the existing entry. If no entry exists, a newentry is written into the queue, and if the number of free entries fallsunder a threshold, the oldest entries are written back to the L2 268 andL3 272 caches.

When a new outermost transaction begins, all existing entries in thestore cache are marked closed so that no new stores can be gathered intothem, and eviction of those entries to L2 268 and L3 272 is started.From that point on, the transactional stores coming out of the LSU 280STQ 260 allocate new entries, or gather into existing transactionalentries. The write-back of those stores into L2 268 and L3 272 isblocked, until the transaction ends successfully; at that pointsubsequent (post-transaction) stores can continue to gather intoexisting entries, until the next transaction closes those entries again.

The store cache 264 is queried on every exclusive or demote XI, andcauses an XI reject if the XI compares to any active entry. If the coreis not completing further instructions while continuously rejecting XIs,the transaction is aborted at a certain threshold to avoid hangs.

The LSU 280 requests a transaction abort when the store cache 264overflows. The LSU 280 detects this condition when it tries to send anew store that cannot merge into an existing entry, and the entire storecache 264 is filled with stores from the current transaction. The storecache 264 is managed as a subset of the L2 268: while transactionallydirty lines can be evicted from the L1 240, they have to stay residentin the L2 268 throughout the transaction. The maximum store footprint isthus limited to the store cache size of 64×128 bytes, and it is alsolimited by the associativity of the L2 268. Since the L2 268 is 8-wayassociative and has 512 rows, it is typically large enough to not causetransaction aborts.

If a transaction aborts, the store cache 264 is notified and all entriesholding transactional data are invalidated. The store cache 264 also hasa mark per doubleword (8 bytes) whether the entry was written by a NTSTGinstruction—those doublewords stay valid across transaction aborts.

Millicode-Implemented Functions

Traditionally, IBM mainframe server processors contain a layer offirmware called millicode which performs complex functions like certainCISC instruction executions, interruption handling, systemsynchronization, and RAS. Millicode includes machine dependentinstructions as well as instructions of the instruction set architecture(ISA) that are fetched and executed from memory similarly toinstructions of application programs and the operating system (OS).Firmware resides in a restricted area of main memory that customerprograms cannot access. When hardware detects a situation that needs toinvoke millicode, the instruction fetching unit 204 switches into“millicode mode” and starts fetching at the appropriate location in themillicode memory area. Millicode may be fetched and executed in the sameway as instructions of the instruction set architecture (ISA), and mayinclude ISA instructions.

For transactional memory, millicode is involved in various complexsituations. Every transaction abort invokes a dedicated millicodesub-routine to perform the necessary abort steps. The transaction-abortmillicode starts by reading special-purpose registers (SPRs) holding thehardware internal abort reason, potential exception reasons, and theaborted instruction address, which millicode then uses to store a TDB ifone is specified. The TBEGIN instruction text is loaded from an SPR toobtain the GR-save-mask, which is needed for millicode to know which GRs238 to restore.

The CPU 114 (FIG. 2) supports a special millicode-only instruction toread out the backup-GRs 224 and copy them into the main GRs 228. TheTBEGIN instruction address is also loaded from an SPR to set the newinstruction address in the PSW to continue execution after the TBEGINonce the millicode abort sub-routine finishes. That PSW may later besaved as program-old PSW in case the abort is caused by a non-filteredprogram interruption.

The TABORT instruction may be millicode implemented; when the IDU 208decodes TABORT, it instructs the instruction fetch unit to branch intoTABORT's millicode, from which millicode branches into the common abortsub-routine.

The Extract Transaction Nesting Depth (ETND) instruction may also bemillicoded, since it is not performance critical; millicode loads thecurrent nesting depth out of a special hardware register and places itinto a GR 228. The PPA instruction is millicoded; it performs theoptimal delay based on the current abort count provided by software asan operand to PPA, and also based on other hardware internal state.

For constrained transactions, millicode may keep track of the number ofaborts. The counter is reset to 0 on successful TEND completion, or ifan interruption into the OS occurs (since it is not known if or when theOS will return to the program). Depending on the current abort count,millicode can invoke certain mechanisms to improve the chance of successfor the subsequent transaction retry. The mechanisms involve, forexample, successively increasing random delays between retries, andreducing the amount of speculative execution to avoid encounteringaborts caused by speculative accesses to data that the transaction isnot actually using. As a last resort, millicode can broadcast to otherCPUs 114 (FIG. 2) to stop all conflicting work, retry the localtransaction, before releasing the other CPUs 114 to continue normalprocessing. Multiple CPUs 114 must be coordinated to not causedeadlocks, so some serialization between millicode instances ondifferent CPUs 114 is required.

Many computer software applications execute in highly parallelenvironments, where multiple concurrently executing programs may allattempt to modify data at a single common memory address. Typically, theaccess to the common memory address may be serialized, for example, byconditional logic, locks, or transactional execution. Transactions mayserialize data access at a finer level than one coarse serialization foran entire database, but may not work well for highly parallel sharing ofdata at a single common memory address. The more accesses made to thecommon memory address in a highly parallel environment, the slower theexecution of the application due to wait time for conditional logic andlocks, and due to transactional aborts for conflicts of transactionalmemory. In one or more embodiments of the present disclosure, computersoftware applications, running in a highly parallel environment,updating data in a single memory address may advantageously utilize anaddressable accumulator to access data at the single memory address. Theaddressable accumulator may atomically access and update data in thememory address replacing the need for conditional logic, locks andtransactional execution and may allow for highly parallel, atomicupdating of data in a single memory address without memory conflict.

An exemplary application advantageously utilizing an addressableaccumulator may be a banking application that maintains a customer'sbank account. The bank account may be updated simultaneously by multiplesources including ATMs, direct deposits, and account debits by amerchant. The banking application may utilize an addressable accumulatorby assigning a customer's bank account to an address that is memorymapped to an addressable accumulator and issuing accumulation requests,related to the bank account, to that address. The addressableaccumulator may recognize the request, on a memory bus, the requestconfigured to have the addressable accumulator cause an update thecustomer's bank account value. The request may signal the ArithmeticLogic Unit to perform the requested update. In an embodiment, theaddressable accumulator may be memory mapped and addressable usingmemory addressing, for example PCIe DMA protocol.

Debiting a customer's bank account in a transactional memory (TM)environment may require fetching the current value of the account frommemory, subtracting the debit amount from the current value and storingthe updated account value to memory. If another transaction, executingon the same bank account, debited the account prior to the firsttransaction storing the updated account value, the transaction may beaborted and restarted. In an addressable accumulator embodiment, a TXtransaction is not needed, instead multiple requests to the addressableaccumulator may debit the account without transactional memory conflictssince the data in the accumulator memory may be accessed directly by thememory subsystem as opposed to data in transactional memory which may beaccessed through the processor 114 (FIG. 10) and cached. Addressableaccumulators may enhance highly parallel execution performance whenmultiple applications simultaneously update data in a common addresslocation by eliminating lost execution time due to transaction aborts aswell as avoiding memory translation and memory caching overhead.

It should be noted that the data in a memory address mapped by anaddressable accumulator may be updated directly by the memory subsystemand may not be cached, in other words, the processor directly accessesthe addressable accumulator rather than by way of a cache. Preferably,an addressable accumulator is compatible with transactional executionbut since the cache may not be utilized, it should also be noted thataccesses to the data in the addressable accumulator may not causetransactional memory conflicts. In a transactional executiontransaction, if the data values in an accumulator address are updatedwithin a transactional execution and the transaction aborts, any TXupdates made to the data values may require computer softwareapplications to provide software recovery to undo any accumulationrequests made during the transaction.

In addition to avoiding wasted CPU cycles on transaction aborts, theaddressable accumulator may improve application performance by avoidingCPU cycles associated with the cache hierarchy. Typically, in a multipleprocessor environment with multiple accumulation requests executing oncommon data, each processor must request the cache line for the commondata exclusively. This may involve a request up the cache hierarchy ofthe processor requesting exclusivity and down the cache hierarchy of theprocessor that owns the requested cache line. The processors may becontinually passing the cache line back and forth in a highly parallelenvironment where multiple processors attempt accumulation on the samedata. Currently, the passing of cache lines may utilize 20 to 50 CPUcycles of execution time, on average. In an embodiment of thedisclosure, the common data may be maintained locally in the addressableaccumulator and may switch between the multiple processor requests everycycle, allowing for higher throughput, less wasted CPU cycles and betterapplication performance than even a Load and Add atomic memory addoperation.

In an embodiment, accumulation requests may be handled, by theaddressable accumulator immediately, or if the addressable accumulatoris busy, enqueued for asynchronous handling. The addressable accumulatormay enqueue multiple concurrent accumulation requests for the sameaddressable accumulator for later processing, allowing a computersoftware application making the accumulation request to progress onother program computations, rather than abort due to memory conflicts,as in a TX implementation. Thus, a processor may send an accumulationrequest to the addressable accumulator and continue processing, eventhough the accumulation request has not yet been fulfilled, but merelyenqueued. The enqueued accumulation request may be fulfilled when theaddressable accumulator has processed all earlier enqueued requests. Inan embodiment, the processor may be notified when the accumulationrequest has completed. In an embodiment, the processor sending theaccumulation request may continue execution without waiting for therequested operation to be completed. Exemplary completion notificationsmay include, but are not limited to, an end of instruction signal with atag identifying the ended accumulation request. The computer softwareapplication may wait for an accumulation requested operation to completeor may continue processing, to be notified asynchronously of theaccumulation request operation's completion. Each accumulation requestoperation executes atomically.

Additionally, the addressable accumulator may free a computer softwareapplication from the need to read a value from memory prior to updatingthe value by replacing the code sequence of load/update/store with asingle accumulation request such as “add 10”. In an embodiment, theaccumulation request may be implemented as one or more new instructionopcodes guaranteed to execute atomically. Each accumulation request mayinclude, but is not limited to, an operation to be executed and thevalue to be used in the operation.

In an embodiment, the addressable accumulator may perform one atomicoperation on the value in the memory address for each accumulationrequest. In another embodiment, the addressable accumulator may have theability to perform multiple operations atomically for an accumulationrequest. The addressable accumulator may keep a history including, butnot limited to, any previous accumulation requests performed, previousvalues of the memory address, time/date stamp of previous accumulationrequests performed and CPU ids and request tags of previous accumulationrequests performed. In an embodiment, accumulation requests may beimplemented to read the addressable accumulator history information andplace it in main memory.

Referring now to FIGS. 4 and 5, an exemplary schematic block diagram 400illustrating an embodiment of a memory mapped addressable accumulator450, hereinafter “accumulator”, attached to a memory bus 826 (FIG. 10)and an interconnect 122 (FIG. 1) in a multiprocessor environment, inaccordance with embodiments of the disclosure. In an embodiment, theaccumulator 450 may be directly attached to a processor memory bus 826(FIG. 10), and directly addressed as if it were a portion of systemmemory. The memory address for accessing the accumulator 450 willhereinafter be referred to as the “accumulator address.” The accumulator450 may act as a direct memory access (DMA) device and may transfer readand write data between main memory 822 (FIG. 10) and the accumulator 450without passing it through the processor 114 (FIG. 10). The accumulator450 may be accessed directly as a location of a memory subsystem and notthrough a cache of a cache-subsystem. The memory subsystem may controlthe actual requests sent on the memory bus 826 (FIG. 10) to theaccumulator 450, arbitrating among multiple processors 114 (FIG. 10)simultaneously sending accumulation requests to same accumulator 450,constraining the accumulation requests to only one at a time. Existingmemory bus 826 (FIG. 10) arbitration schemes may include, but are notlimited to, a round robin list of requestors and multiple memory bussesper CPU. For the memory bus arbitration scheme of a single requestor percycle, as in the round robin arbitration, the accumulator 450 maycapture the requests in the order they were received and make theupdates atomically even if the update requires multiple cycles tocomplete. For the memory bus arbitration scheme of multiple memorybusses per processor 114 (FIG. 10) potentially writing to the sameaccumulator 450 at the same time, the accumulator 450 may have multipleports, one port having priority over the others.

The memory bus 826 (FIG. 10) may be bidirectional or unidirectional andmay transfer data including, but not limited to, an address and bitsdenoting which bytes in a data partition may be involved in the read orwrite operation. Typically, data transferred on the memory bus 826 (FIG.10) are 8 or 16 bytes wide. The accumulator 450 may monitor the memorybus 826 (FIG. 10) and detect and operate on any read or write requestson the memory bus 826 (FIG. 10) with an address corresponding to theaccumulator address.

The accumulator 450 may additionally be attached to an interconnect 122(FIG. 1) to communicate instruction completions and threshold alertswith the processors 114 a, 114 b . . . 114 n in the computer system 1600(FIG. 10). Multiple processors 114 a, 114 b . . . 114 n may all sendaccumulation requests to the same accumulator 450.

In another embodiment, the accumulator 450 may be connected with PCIexpress (PCIe) rather than a memory bus 826 (FIG. 10). PCIecommunication is encapsulated in packets. A transaction layer may beimplemented in the accumulator 450 to handle the packeting andde-packeting of data and status message traffic. The transaction layermay provide a well defined interface where PCIe packetized data may beprepared and provided or received from the PCIe connection. Thetransaction layer may receive read and write requests and may createrequest packets for transmission to a link layer. All requests may beimplemented as split transactions and some of the request packets mayrequire a response packet. Each packet may have a unique identifier thatenables response packets to be directed to the correct originator.

The transaction layer may provide four address spaces—three PCI addressspaces (memory, I/O, and configuration) and message space. PCI 2.2introduced an alternate method of propagating system interrupts calledmessage signaled interrupt (MSI). Here a special-format memory-writetransaction may be used instead of a hard-wired sideband signal as anoptional capability in a PCI 2.2 system. The PCIe specification reusesthe MSI concept as a primary method for interrupt processing and uses amessage space to accept all prior sideband signals, such as interrupts,power-management requests, and resets, as in-band messages. Other“special cycles” within the PCI 2.2 specification, such as interruptacknowledge, are also implemented as in-band messages. PCIe messages maybe thought of as “virtual wires” because their effect is to eliminatethe wide array of sideband signals currently used in a platformimplementation.

With continuing reference to FIG. 5, an exemplary schematic blockdiagram 500 illustrates a memory mapped addressable accumulator 450 andan input accumulation request data packet 440, in accordance withembodiments of the disclosure. In an embodiment, the accumulator 450 mayinclude an accumulator interface 420, a controller 430, an arithmeticlogic unit 470, an accumulator queuer 410 with its queue 415, amiscellaneous register 451, a threshold register 452, an accumulationregister 453 and an address register 460. The accumulator 450 may beattached by a memory or system bus 826 (FIG. 10) and an interconnect 122(FIG. 1) to one or more processors 114 (FIG. 10). The accumulatoraddress may be maintained in the address register 460.

In an embodiment, the accumulator interface 420 may act as thecommunication interface with the bus 826 (FIG. 10) and the interconnect122 (FIG. 1). The accumulator interface 420 may monitor the bus 826(FIG. 10) and may determine if the data on the bus 826 (FIG. 10) ismeant for this accumulator 450 by comparing the address, M, of any dataon the bus 826 (FIG. 10) with the address in one or more addressregisters 460. The address in address register 460 may be a singleaddress, multiple addresses or a range of addresses supported by theaccumulator 450. The address register 460 may be a programmable registerinitialized by an accumulation request executed by a processor 114 (FIG.10) with direct control of the accumulator 450 or may be an addressestablished at system initialization. The accumulator interface 420 maycommunicate with the accumulator queuer 410. Communications may include,but are not limited to the accumulator interface 420 passing anaccumulation request data packet 440 to the accumulator queuer 410 to beadded as an entry on its queue 415 and receiving, from the accumulatorqueuer 410, a notification if a queue length threshold has been reached.The accumulator interface 420 may additionally communicate with thecontroller 430. Communications may include, but are not limited to, theaccumulator interface 420 indicating, to the controller 430, that anaccumulation request data packet 440 has been received, that anaccumulation request data packet 440 has been added to the queue 415 andreceiving, from the controller 430, a notification that the arithmeticlogic unit 470 is currently active, that an accumulation request hascompleted or a threshold has been reached. The accumulator interface 420may signal all processors 114 (FIG. 10) through the interconnect 122(FIG. 1) that the accumulation request has completed (the requestedoperation has been performed) and, if applicable, that a threshold hasbeen reached. The accumulator interface 420 may additionally place dataon the bus 826 (FIG. 10) in response to the completion of anaccumulation request instruction 444 requesting a “retrieve-value” ofthe accumulated data value in the accumulation register 453. Additionalinformation from the miscellaneous register 451 may also be added to thedata placed on the bus 826 (FIG. 10).

In an embodiment, the controller 430 may monitor the number ofaccumulation request data packet 440 a, 440 b, 440 c entries queued onthe queue 415 to know when all queued accumulation request data packets440 a, 440 b, 440 c have been processed. The controller 430 may alsomonitor the activity of the arithmetic logic unit 470. When thecontroller 430 finds the arithmetic logic unit 470 available, thecontroller 430 may communicate with the accumulator queuer 410 to havean entry dequeued from its queue 415. The controller 430 may receive theinstruction 444 portion of the dequeued entry while the arithmetic logicunit 470 may receive the value 446 portion of the entry. The controller430 may decode the instruction 444 portion of the entry and pass thedecoded arithmetic operation to the arithmetic logic unit 470. If theinstruction 444 portion of the entry is a complex operation, thecontroller 430 may decode the function into multiple simple arithmeticoperations and may pass each simplified arithmetic operation, in turn,to the arithmetic logic unit 470. The controller may also store anyadditional information included in the instruction 444 portion of theentry into the miscellaneous register 451. The controller 430 mayreceive, from the arithmetic logic unit 470, a notification that theinstruction has completed and, if applicable, a notification that athreshold value has been reached. The controller 430 may send thecompletion notification and, if applicable, the threshold notificationto the accumulator interface 420. The controller 430 may also store aninitial value into the threshold register 452 for a“store-threshold-value” accumulation request instruction 444. Thethreshold value set may indicate whether it is a low threshold value ora high threshold value. In another embodiment, there may be multiplethreshold registers 453 allowing for both a low threshold value and ahigh threshold value.

In an embodiment, the accumulator queuer 410 may maintain a queue 415 ofaccumulation request data packet 440 a, 440 b, 440 c entries. Theaccumulator queuer 410 may receive an accumulation request data packet440 from the accumulator interface 420 and add it to the queue 415 as aqueue entry. The accumulator queuer 410 may determine if a queue lengththreshold has been reached and if so, may notify the accumulatorinterface 420. In an embodiment, the queue length threshold may be basedon the size of the queue 415. In another embodiment, a programmablequeue threshold register may be utilized to allow for dynamicmodification of the queue length threshold. The accumulator queuer 410may also receive a request from the controller 430 to dequeue an entryfrom the queue 415. The accumulator queuer 410 may dequeue the entryfrom the queue 415 in any order, including but not limited to,First-in-First-Out, application priority, a priority supplied with theaccumulation request and Last-in-First-Out and pass the instruction 444portion of the dequeued entry to the controller 430 and pass the value446 portion of the dequeued entry to the arithmetic logic unit 470. Thequeue 415 may be either a fixed size or grow and shrink dynamically asentries are enqueued and dequeued.

In an embodiment, the arithmetic logic unit 470 may receive a decodedarithmetic operation from the controller 430 and a data value 446 fromthe accumulator dequeuer 410. The arithmetic logic unit 470 may executethe arithmetic operation utilizing the value in the accumulationregister 453 and the received data value 446. The arithmetic logic unit470 may store the results of the executed arithmetic operation into theaccumulation register 453. For a decoded operation of a“store-initial-value” instruction 444, the arithmetic logic unit 470 maystore the data value 446 into the accumulation register 453 withoutperforming any other arithmetic operations on the data. The arithmeticlogic unit 470 may additionally compare the resulting value of theexecuted operation with the value in the threshold register 452. For aresulting value that has reached or passed the threshold value in thethreshold register 452, the arithmetic logic unit 470 may notify thecontroller 430 that the threshold has been reached.

In an embodiment, the miscellaneous register 451 may include anyinformation passed in the accumulation request data packet 440 with theinstruction 444. The controller 430 may store the information in themiscellaneous register 451 for each instruction 444 decoded. Themiscellaneous register 451 value may be sent to the accumulatorinterface 420 with each completion and threshold notification. Inanother embodiment, the miscellaneous register 451 value may be storedto accumulator memory or non-accumulator memory for each accumulationrequest in order to create a history of the accumulator 450 that may beloaded at a later time by the processor 114 (FIG. 10).

An exemplary accumulation request data packet 440 for address M isshown. The accumulation request data packet 440 may be received from thememory bus 826 (FIG. 10) by the accumulator 450 that is memory mapped toaddress M. The accumulation request data packet 440 may include, but isnot limited to, an accumulation instruction 444 and an immediate datavalue 446 to be accumulated. The accumulation instruction 444 mayinclude, but is not limited to, an operation type, a data type for thevalue 446, an initiating CPU identifier, a time and date stamp and anaccumulation request identifier tag. In an embodiment, the processor 114(FIG. 10) may be configured such that a computer software applicationmay issue a store instruction which includes an accumulation requestdata packet 440 and an accumulator address M. The processor's 114 (FIG.10) store instruction may recognize that the address M is not in thecache but out on the memory bus 826 (FIG. 10) and the memory subsystemmay recognize the address as a direct memory attached (DMA) device andadd the accumulation request data packet 440 to the bus 826 (FIG. 10).In other embodiments the accumulation request data packet 440 may bememory data or a PCIe packet.

In an embodiment, the accumulator 450 may be a discreet memory mappeddevice, attached to the computer system 1600 (FIG. 10). Computer system1600 (FIG. 10) may attach multiple accumulator 450 devicessimultaneously, each with its own accumulator address, accumulatorinterface 420, controller 430, accumulator queuer 410 and its queue 415,arithmetic logic unit 470, address register 460, miscellaneous register451, threshold register 452 and accumulation register 453. In anembodiment, a single accumulator 450 may support multiple accumulatoraddresses, accumulator queues 415, address registers 460, miscellaneousregisters 451, threshold registers 452 and accumulation registers 453.

In another embodiment, the accumulator 450 may be implemented on a fieldprogrammable gate array (FPGA). The FPGA may be attached to memory by abus 826 (FIG. 10) and may be memory mapped. Computer system 1600 (FIG.10) may attach multiple accumulators 450 implemented on FPGAssimultaneously, each with its own each accumulator address, accumulatorinterface 420, controller 430, accumulator queuer 410 and its queue 415,arithmetic logic unit 470, address register 460, miscellaneous register451, threshold register 452 and accumulation register 453. Computersystem 1600 (FIG. 10) may attach multiple accumulators 450 implementedon FPGAs simultaneously with multiple accumulator 450 devices. In anembodiment, a single accumulator 450 implemented on an FPGA may supportmultiple accumulator addresses, accumulator queues 415, addressregisters 460, miscellaneous registers 451, threshold registers 452 andaccumulation registers 453.

In another embodiment, described in more detail with reference to FIGS.7 and 8, the processor 114 (FIG. 10) may emulate an accumulator 450. Theaccumulator addresses may be system-defined or may be dynamicallydefined by computer software applications. The processor 114 (FIG. 10)may monitor the buses 826 (FIG. 10) for activity to the accumulatoraddresses. The processor 114 (FIG. 10) may simultaneously emulatemultiple accumulators 750 (FIG. 7) for multiple accumulator addresses.Computer system 1600 (FIG. 10) may emulate multiple accumulators 750(FIG. 7) simultaneously with multiple accumulator 450 devices as well asmultiple accumulators 450 implemented on FPGAs.

Now referring to FIG. 6 and with continuing reference to FIG. 5, aflowchart 600 illustrates atomic accumulation of a value in accumulatormemory by an accumulator 450, illustrated within the data processingenvironment of FIG. 10. The flowchart 600 illustrates an embodiment inwhich the accumulator 450 atomically updates the value in accumulatormemory. With continued reference to the banking application exampleabove, a customer “A” may wish to open a bank account and deposit$750.33 into the account. The banking application may place the bankaccount into an exemplary accumulator address 1500. The accumulator 450may be set to respond to any memory bus write or read to location 1500.To initialize the value of the account for customer “A”, the bankingapplication, running on a processor 114 (FIG. 10), may create anaccumulation request data packet 440 that includes a“store-initial-value” instruction 444 in the first 8 bytes, and anencoded value $750.33 in the second 8 bytes. The format of the value 446may include, but is not limited to, a format denoted by the instructiontype such as a Binary Coded Decimal number with a scale of 10 to thepower −2 and a hexadecimal number “00000000 00075033”. The processor 114(FIG. 10) may issue a store for the accumulation request data packet 440representing the “store-initial-value” instruction 444 and value 446 formemory address 1500. The write operation may be placed on the memorybus, at 505, with the 16 byte accumulation request data packet 440. Theaccumulation request data packet 440 may also include an instructiontag, to be sent back when the instruction completes, to uniquelyidentify the processor 114 (FIG. 10) and the instruction 444. Exemplaryaccumulation request instructions may include, but are not limited to,“store-initial-value”—to store an initial value into the accumulationregister 453, “store-threshold-value”—to store a threshold value intothe threshold register 452, “add-value”—to add the received value 446 tothe value in the accumulation register 453, “subtract-value”—to subtractthe received value 446 from the value in the accumulator register 453,“function” request—to perform a function (such as: add a percentage) onthe value in the accumulation register 453, and “retrieve-value”—to copythe value from the accumulation register 453 to another location. Theother location receiving the value from the accumulation register 453may include, but is not limited to, a system memory location, a systemcontrol area, a register, or a computer software application providedlocation. The accumulator 450 with the accumulator address 1500 (in thebanking example) may receive the accumulation request data packet 440from the memory bus 826 (FIG. 10), at 510.

If the accumulator 450 is currently busy updating the accumulationregister 453, as determined at 515, when the accumulation request datapacket 440 is received, the accumulator may, at 530, queue the receivedaccumulation request data packet 440 on the queue 415 to be processedlater. If the queue 415 reaches a full threshold, as determined at 535,an interrupt may be issued, at 540, to alert the processors 114 (FIG.10) to delay sending any new accumulation requests for this accumulatoraddress until the queue 415 clears. In an embodiment, a full queue 415may cause the accumulation request to stall, or may cause it to berejected. In response to the queue threshold interrupt, the processors114 (FIG. 10) may delay sending new accumulation requests to theaccumulator address issuing the interrupt based on a time delay or basedon receiving an additional interrupt signaling that the queue 415 hascleared. If the accumulator 450 continues to be busy updating theaccumulation register 453 after the received accumulation request datapacket 440 is queued, as determined at 545, the accumulation requestdata packet 440 may remain queued until any other accumulation datapackets 440 a, 440 b, 440 c on the queue 415 have been processed.

If the accumulator 450 is not currently updating the accumulationregister 453 when the accumulation request data packet 440 is received,as determined at 515, the accumulator may update, at 520, the value inthe accumulation register 453 as requested in the received accumulationrequest data packet 440. For the exemplary bank account initializationrequest, the “store-initial-value” instruction 444 may be decoded andthe initialization performed, at 520, by the arithmetic logic unit 470.For the simple store-initial-value operation, the value 446 may bepassed through the arithmetic logic unit 470 and the accumulationregister 453 updated to the value $750.33.

Once the accumulator 450 updates the value in the accumulation register453, at 520, the value may be checked, at 555, against a data thresholdvalue. For an accumulation register 453 value either below a minimumthreshold or above a maximum threshold, the accumulator 450 may, at 560,issue an interrupt to alert the processors 114 (FIG. 10) that the datathreshold value may have been reached. At 570, the accumulator 450 maysignal the completion of the accumulation request, identifying theinitiating processor 114 (FIG. 10) and the completed accumulationrequest. A completed accumulation request may allow the computersoftware application to continue processing any program instructionsdependent on the updated value in the accumulator memory.

If the accumulator 450 is now free, either after queuing the receivedaccumulation request data packet 440 onto the queue 415 or aftercompleting a prior accumulation request, determined at 545, and thereare queued accumulation request data packets 440 a, 440 b, 440 c on thequeue 415, as determined at 547, the accumulator 450 may, at 550,dequeue an accumulation request data packet 440 a, 440 b, 440 c from thequeue 415 and update, at 520, the value in the accumulation register 453as requested in the dequeued accumulation request data packet 440 a, 440b, 440 c.

Continuing the banking application example, the banking application hasbeen notified that the account is initialized. Customer “A” may now stopat an ATM and withdraw $20.00. The ATM may be connected to a computersystem 1600 (FIG. 10) in which the processor 114 (FIG. 10) may create anaccumulation request data packet 440 which includes an instruction 444to “subtract-value” and a value 446 set to the hexadecimal number“00000000 00002000”. The accumulation request data packet 440 may besent to customer “A”s bank account at accumulator address 1500. Theaccumulator 450 monitoring address 1500 may receive the accumulationrequest data packet 440 for the ATM withdrawal and either process itimmediately or enqueue it on the queue 415 for later processing. Whenthe accumulation request data packet 440 for the ATM withdrawal isprocessed, the accumulator 450 may signal the arithmetic logic unit 470to subtract the value $20.00 from the current value in the accumulationregister 453. The current value of $750.33 would be reduced to $730.33.The value in the threshold register 452 may be checked and may triggeran alert if the new balance in the account is below the minimum valueallowed for withdrawal at an ATM. The threshold alert may signal allprocessors 114 (FIG. 10) that customer “A”s bank account balance isdangerously low.

In another embodiment, all received accumulation request data packets440 may be initially enqueued on the queue 415 for later processing. Inan embodiment, the enqueued accumulation request data packets 440 a, 440b, 440 c may be sorted and prioritized prior to being dequeued from thequeue 415 for processing.

Referring now to FIG. 7, an exemplary schematic block diagram 700illustrating processor 114 (FIG. 10) emulating a memory mappedaddressable accumulator 450 (FIG. 5), in accordance with embodiments ofthe disclosure. In an embodiment, the emulated accumulator 750 may beemulated by the processor 114 (FIG. 10) with any mixture of hardware,software and microcode. In an embodiment, the accumulator addressregister 460 (FIG. 5) and controller 730 may be hardware componentsrunning concurrently with other processor 114 (FIG. 10) functions,continually monitoring the memory bus 826 (FIG. 10) for accumulationrequests. An emulated threshold register 752, emulated accumulationregister 753, emulated miscellaneous register 751 and emulated queue 715may all reside in the processor's 114 (FIG. 10) local memory 710. In theembodiment shown, the accumulator address register 460 (FIG. 5) mayrecognize an accumulation request on the memory bus 826 (FIG. 10) forthis emulated accumulator 750 and send it to the controller 730. Thecontroller may maintain pointers 731, 732, 733 to the emulatedaccumulator 750 registers 751, 752, 753. The controller 730, in thisembodiment, may both load and store any data values into the emulatedregisters 751, 752, 753 as well as manage the emulated queue 715. In theembodiment shown, the emulated queue 715 is a fixed length, circularqueue with the controller 730 maintaining the circular queue pointers734, 735, 736, and 737. The controller 730 may additionally maintain thecount of emulated queue 715 entries, to prevent overwriting any entriesand to signal, over the interconnect 122 (FIG. 1), queue thresholdalerts to all processors 114 (FIG. 10). The controller 730 may alsocommunicate with the arithmetic logic unit 470 (FIG. 5) to perform theaccumulation requests, to check the results against the threshold valuesin the emulated threshold register 752 and to signal alerts for anythreshold values reached during the accumulation.

In another embodiment, the controller 730 may be a software routine andthe emulated queue 715 may be managed as any one of a fixed lengthqueue, variable length queue or a dynamic link list queue.

In an embodiment, the emulated accumulator 750 may have an accumulatoraddress determined by a computer software application and may bedynamic. The computer software application may signal to the processor114 (FIG. 10) to begin emulation on a specific memory address. Theprocessor 114 (FIG. 10) may set the accumulator address register 460(FIG. 5) to the specified memory address and start the emulatedaccumulator 750 monitoring the bus 826 (FIG. 10) for the specifiedmemory address. When the computer software application completes, it maysignal the processor 114 (FIG. 10) to stop monitoring the specifiedmemory address, stop the emulated accumulator 750 for the specifiedmemory address and reset the specified memory address as anon-accumulator memory address.

During emulation, processor microcode may perform the emulatedaccumulator 750 functions using host memory space unavailable to theoperating system or computer software applications. In another emulationenvironment, a run-time program may monitor memory accesses and cause aninterruption, trap or call to a software routine that emulates theaccumulator function. In another embodiment, a compiler may insert callsto the software routine that emulates the accumulator function.

Referring now to FIG. 8 and with continuing reference to FIG. 7, anexemplary schematic block diagram 1000 illustrates processor 114 (FIG.10) emulating multiple memory mapped addressable accumulators 750 a-750n, in accordance with embodiments of the disclosure. In an embodiment,the processor 114 (FIG. 10) may emulate each accumulator 750 a-750 nwith any mixture of hardware, software and microcode and each emulatedaccumulator 750 a-750 n may include, but is not limited to, anaccumulator address register 460 (FIG. 5), a controller 730, an emulatedthreshold register 752, an emulated accumulation register 753, anemulated miscellaneous register 751, an emulated queue 715, and anarithmetic logic unit 470 (FIG. 5). Each emulated accumulator 750 a-750n may reside in processor 114 (FIG. 10) local memory. Each emulatedaccumulator 750 a-750 n may monitor the memory bus 826 (FIG. 10) foraccumulation requests directed to the memory address in its addressregister 460 (FIG. 5). Each emulated accumulator 750 a-750 n may signalinstruction completion and alerts to all processors 114 (FIG. 10)through the interconnect 122 (FIG. 1). Emulating large numbers ofaccumulators 750 a-750 n simultaneously may advantageously replacetransactional execution for computer software applications that executein highly parallel environments due to the atomicity of memory updatesby the emulated accumulators 750 a-750 n and the lack of memoryconflicts for accesses to data in accumulator addresses.

Referring now to FIG. 9 and with continuing reference to FIG. 5, anexemplary schematic block diagram 1100 illustrates one accumulator 450mapping multiple virtual accumulators, each with a discreet accumulatoraddress, in accordance with embodiments of the disclosure. The physicalaccumulator 450, mapping multiple virtual accumulators, may be adiscreet device, an FPGA or a processor 114 (FIG. 10) emulatedaccumulator 750. In an embodiment, there may be only one interface tothe memory bus 826 (FIG. 10) and interconnect 122 (FIG. 1). In anembodiment, the physical accumulator 450, 750 may include multipleaccumulator address registers 460 (one for each virtual accumulatoraddress being monitored by the physical accumulator 450, 750), multipleaccumulation registers 453 (one for each virtual accumulator to hold theaccumulation value for the associated virtual accumulator), multiplethreshold registers 452 (one or more for each virtual accumulator tohold the data threshold values for the associated virtual accumulator),multiple miscellaneous registers 451 (one or more for each virtualaccumulator to hold accumulation request information of requests for theassociated virtual accumulator) and an accumulator queuer 410 and queue415 for each virtual accumulator. In an embodiment, one threshold valuemay be shared among all virtual accumulators mapped by the physicalaccumulator 450, 750. In an embodiment one accumulator interface 410,one controller 430 and one arithmetic logic unit 470 may handle all thevirtual accumulators mapped by the physical accumulator 450, 750. Theaccumulator interface 410 may recognize the one or more memory addressesstored in the address registers 460 and may queue any accumulationrequest to the appropriate queue 415. The arithmetic logic unit 470 maybe time multiplexed among the multiple accumulation operations utilizingthe queuing structure of the requests. Including only one arithmeticlogic unit 470 for accumulations and threshold compares in a physicalaccumulator 450, 750 mapping multiple virtual accumulators may moreefficiently utilize hardware components.

Referring now to FIG. 10, computer system 1600 may include respectivesets of internal components 800 and external components 900. Each of thesets of internal components 800 includes one or more processors 114; oneor more computer-readable RAMs 822; one or more computer-readable ROMs824 on one or more buses 826; one or more operating systems 828; one ormore software applications 829; and one or more computer-readabletangible storage devices 830. The one or more operating systems 828 arestored on one or more of the respective computer-readable tangiblestorage devices 830 for execution by one or more of the respectiveprocessors 114 via one or more of the respective RAMs 822 (whichtypically include cache memory). The computer system 1600, in anembodiment of the disclosure, supports one or more accumulator addressesand one or more memory mapped addressable accumulators 450. Theprocessors 114, in this embodiment, are also configured to identifyaccumulation requests and the one or more accumulator addresses. In theembodiment illustrated in FIG. 10, each of the computer-readabletangible storage devices 830 is a magnetic disk storage device of aninternal hard drive. Alternatively, each of the computer-readabletangible storage devices 830 is a semiconductor storage device such asROM 824, EPROM, flash memory or any other computer-readable tangiblestorage device that can store a computer program and digitalinformation.

Each set of internal components 800 also includes a R/W drive orinterface 832 to read from and write to one or more computer-readabletangible storage devices 936 such as a CD-ROM, DVD, SSD, memory stick,magnetic tape, magnetic disk, optical disk or semiconductor storagedevice.

Each set of internal components 800 may also include network adapters(or switch port cards) or interfaces 836 such as a TCP/IP adapter cards,wireless WI-FI interface cards, or 3G or 4G wireless interface cards orother wired or wireless communication links. The firmware 838 andoperating system 828 that are associated with computer system 1600, canbe downloaded to computer system 1600 from an external computer (e.g.,server) via a network (for example, the Internet, a local area networkor other, wide area network) and respective network adapters orinterfaces 836. From the network adapters (or switch port adapters) orinterfaces 836, the firmware 838 and operating system 828 associatedwith computer system 1600 are loaded into the respective hard drive 830and network adapter 836. The network may comprise copper wires, opticalfibers, wireless transmission, routers, firewalls, switches, gatewaycomputers and/or edge servers.

Each of the sets of external components 900 can include a computerdisplay monitor 920, a keyboard 930, and a computer mouse 934. Externalcomponents 900 can also include touch screens, virtual keyboards, touchpads, pointing devices, and other human interface devices. Each of thesets of internal components 800 also includes device drivers 840 tointerface to computer display monitor 920, keyboard 930, and computermouse 934. The device drivers 840, R/W drive or interface 832 andnetwork adapter or interface 836 comprise hardware and software (storedin storage device 830 and/or ROM 824).

Referring now to FIGS. 11 and 12, in an embodiment of the disclosure,memory updates may be accumulated atomically in a computer systemconfigured with an accumulator that is memory mapped. In an embodiment,the accumulator may have an accumulator memory and an accumulator queueand may be configured to communicatively couple to a processor. In anembodiment, the memory mapped accumulator may be emulated by theprocessor, the processor's local memory or registers allocated for anemulated accumulator address, an emulated accumulator value, and anemulated accumulator instruction queue and an emulated threshold value.The processor's arithmetic logic unit may be used as the accumulatorarithmetic logic unit and may include an adder and may further comparethe emulated threshold value with the emulated accumulator value todetermine if the value threshold is reached. In an embodiment, theprocessor may emulate a plurality of accumulators, the plurality ofemulated accumulator addresses initialized based on a plurality ofprogram specified memory addresses, each of the plurality of emulatedaccumulator addresses for controlling one of the plurality of emulatedaccumulator values. In an embodiment, the processor may emulate theplurality of arithmetic logic units with a single arithmetic logic unit.In another embodiment, the accumulator may support a plurality ofaccumulator addresses having corresponding accumulator memories andcorresponding accumulator queues and the accumulator may determine, uponreceiving the accumulation request, which accumulator address, of theplurality of accumulator addresses supported, is to be used forprocessing the request.

In an embodiment, the accumulator may, at 1210, receive, from theprocessor, an accumulation request directed to an accumulator address.The accumulation request includes an accumulator operation identifierand data. The accumulator may, at 1220, determine whether it canimmediately process the request. Based on determining, at 1220, that theaccumulator can immediately process the request, immediately processing,at 1230, the request. Processing the request includes atomicallyupdating a value in the accumulator memory based on the operationidentifier and data of the accumulation request. In an embodiment, basedon the accumulator operation being a store-initial-value operation,determined at 1240, storing an initial value into the accumulatormemory. In an embodiment, based on the accumulator operation being aretrieve-value operation, determined at 1250, retrieving the value fromthe accumulator memory and storing the value to a non-accumulatorlocation. Based on the value resulting from processing the requestreaching a threshold value, determined at 1260, the accumulator maysignal the processor the value threshold being reached. At 1270, theaccumulator may signal the processor the completion of the accumulationrequest. The accumulator may continue processing, at 1220, to process aqueued accumulation request from the accumulator queue or a new requestreceived from the processor.

Based on the accumulator determining, at 1220, that request can not beimmediately processed, based on the accumulator actively processinganother accumulation request, the accumulator may queue, at 1280, theaccumulation request for later processing. Based on reaching a queuelength threshold after queuing the accumulation request for laterprocessing, the accumulator may signal the processor, at 1290, the queuelength threshold being reached. The accumulator may continue processing,at 1220, to process a queued accumulation request from the accumulatorqueue or a new request received from the processor.

Various embodiments of the invention may be implemented in a dataprocessing system suitable for storing and/or executing program codethat includes at least one processor coupled directly or indirectly tomemory elements through a system bus. The memory elements include, forinstance, local memory employed during actual execution of the programcode, bulk storage, and cache memory which provide temporary storage ofat least some program code in order to reduce the number of times codemust be retrieved from bulk storage during execution.

Input/Output or I/O devices (including, but not limited to, keyboards,displays, pointing devices, DASD, tape, CDs, DVDs, thumb drives andother memory media, etc.) can be coupled to the system either directlyor through intervening I/O controllers. Network adapters may also becoupled to the system to enable the data processing system to becomecoupled to other data processing systems or remote printers or storagedevices through intervening private or public networks. Modems, cablemodems, and Ethernet cards are just a few of the available types ofnetwork adapters.

The present invention may be a system, a method, and/or a computerprogram product. The computer program product may include a computerreadable storage medium (or media) having computer readable programinstructions thereon for causing a processor to carry out aspects of thepresent invention.

The computer readable storage medium can be a tangible device that canretain and store instructions for use by an instruction executiondevice. The computer readable storage medium may be, for example, but isnot limited to, an electronic storage device, a magnetic storage device,an optical storage device, an electromagnetic storage device, asemiconductor storage device, or any suitable combination of theforegoing. A non-exhaustive list of more specific examples of thecomputer readable storage medium includes the following: a portablecomputer diskette, a hard disk, a random access memory (RAM), aread-only memory (ROM), an erasable programmable read-only memory (EPROMor Flash memory), a static random access memory (SRAM), a portablecompact disc read-only memory (CD-ROM), a digital versatile disk (DVD),a memory stick, a floppy disk, a mechanically encoded device such aspunch-cards or raised structures in a groove having instructionsrecorded thereon, and any suitable combination of the foregoing. Acomputer readable storage medium, as used herein, is not to be construedas being transitory signals per se, such as radio waves or other freelypropagating electromagnetic waves, electromagnetic waves propagatingthrough a waveguide or other transmission media (e.g., light pulsespassing through a fiber-optic cable), or electrical signals transmittedthrough a wire.

Computer readable program instructions described herein can bedownloaded to respective computing/processing devices from a computerreadable storage medium or to an external computer or external storagedevice via a network, for example, the Internet, a local area network, awide area network and/or a wireless network. The network may comprisecopper transmission cables, optical transmission fibers, wirelesstransmission, routers, firewalls, switches, gateway computers and/oredge servers. A network adapter card or network interface in eachcomputing/processing device receives computer readable programinstructions from the network and forwards the computer readable programinstructions for storage in a computer readable storage medium withinthe respective computing/processing device.

Computer readable program instructions for carrying out operations ofthe present invention may be assembler instructions,instruction-set-architecture (ISA) instructions, machine instructions,machine dependent instructions, microcode, firmware instructions,state-setting data, or either source code or object code written in anycombination of one or more programming languages, including an objectoriented programming language such as Java, Smalltalk, C++ or the like,and conventional procedural programming languages, such as the “C”programming language or similar programming languages. The computerreadable program instructions may execute entirely on the user'scomputer, partly on the user's computer, as a stand-alone softwarepackage, partly on the user's computer and partly on a remote computeror entirely on the remote computer or server. In the latter scenario,the remote computer may be connected to the user's computer through anytype of network, including a local area network (LAN) or a wide areanetwork (WAN), or the connection may be made to an external computer(for example, through the Internet using an Internet Service Provider).In some embodiments, electronic circuitry including, for example,programmable logic circuitry, field-programmable gate arrays (FPGA), orprogrammable logic arrays (PLA) may execute the computer readableprogram instructions by utilizing state information of the computerreadable program instructions to personalize the electronic circuitry,in order to perform aspects of the present invention.

Aspects of the present invention are described herein with reference toflowchart illustrations and/or block diagrams of methods, apparatus(systems), and computer program products according to embodiments of theinvention. It will be understood that each block of the flowchartillustrations and/or block diagrams, and combinations of blocks in theflowchart illustrations and/or block diagrams, can be implemented bycomputer readable program instructions.

These computer readable program instructions may be provided to aprocessor of a general purpose computer, special purpose computer, orother programmable data processing apparatus to produce a machine, suchthat the instructions, which execute via the processor of the computeror other programmable data processing apparatus, create means forimplementing the functions/acts specified in the flowchart and/or blockdiagram block or blocks. These computer readable program instructionsmay also be stored in a computer readable storage medium that can directa computer, a programmable data processing apparatus, and/or otherdevices to function in a particular manner, such that the computerreadable storage medium having instructions stored therein comprises anarticle of manufacture including instructions which implement aspects ofthe function/act specified in the flowchart and/or block diagram blockor blocks.

The computer readable program instructions may also be loaded onto acomputer, other programmable data processing apparatus, or other deviceto cause a series of operational steps to be performed on the computer,other programmable apparatus or other device to produce a computerimplemented process, such that the instructions which execute on thecomputer, other programmable apparatus, or other device implement thefunctions/acts specified in the flowchart and/or block diagram block orblocks.

The flowchart and block diagrams in the Figures illustrate thearchitecture, functionality, and operation of possible implementationsof systems, methods, and computer program products according to variousembodiments of the present invention. In this regard, each block in theflowchart or block diagrams may represent a module, segment, or portionof instructions, which comprises one or more executable instructions forimplementing the specified logical function(s). In some alternativeimplementations, the functions noted in the block may occur out of theorder noted in the figures. For example, two blocks shown in successionmay, in fact, be executed substantially concurrently, or the blocks maysometimes be executed in the reverse order, depending upon thefunctionality involved. It will also be noted that each block of theblock diagrams and/or flowchart illustration, and combinations of blocksin the block diagrams and/or flowchart illustration, can be implementedby special purpose hardware-based systems that perform the specifiedfunctions or acts or carry out combinations of special purpose hardwareand computer instructions.

Although one or more examples have been provided herein, these are onlyexamples. Many variations are possible without departing from the spiritof the present invention. For instance, processing environments otherthan the examples provided herein may include and/or benefit from one ormore aspects of the present invention. Further, the environment need notbe based on the z/Architecture®, but instead can be based on otherarchitectures offered by, for instance, IBM®, Intel®, Sun Microsystems,as well as others. Yet further, the environment can include multipleprocessors, be partitioned, and/or be coupled to other systems, asexamples.

As used herein, the term “obtaining” includes, but is not limited to,fetching, receiving, having, providing, being provided, creating,developing, etc.

The flow diagrams depicted herein are just examples. There may be manyvariations to these diagrams or the steps (or operations) describedtherein without departing from the spirit of the invention. Forinstance, the steps may be performed in a differing order, or steps maybe added, deleted, or modified. All of these variations are considered apart of the claimed invention.

Although preferred embodiments have been depicted and described indetail herein, it will be apparent to those skilled in the relevant artthat various modifications, additions, substitutions and the like can bemade without departing from the spirit of the invention, and these are,therefore, considered to be within the scope of the invention, asdefined in the following claims.

What is claimed is:
 1. A computer implemented method for atomicallyaccumulating memory updates in a computer system configured with anaccumulator that is memory mapped, the accumulator having an accumulatormemory and an accumulator queue, the accumulator configured tocommunicatively couple to a processor wherein the accumulator isconfigured to perform a method, the method comprising: receiving fromthe processor, by the accumulator, an accumulation request directed toan accumulator address, the accumulation request comprising anaccumulator operation identifier and data; based on determining, by theaccumulator, that the accumulator can immediately process the request,immediately processing the request, the processing comprising atomicallyupdating a value in the accumulator memory, by the accumulator, based onthe operation identifier and data of the accumulation request; and basedon determining, by the accumulator, that the accumulator is activelyprocessing another accumulation request, queuing, by the accumulator,the accumulation request for later processing; and signaling theprocessor, by the accumulator, a completion of the accumulation request.2. The method according to claim 1, further comprising: based on theaccumulator operation being a store-initial-value operation, storing aninitial value into the accumulator memory; and based on the accumulatoroperation being a retrieve-value operation, performing a) and b)comprising: a) retrieving the value from the accumulator memory; and b)storing the value to a non-accumulator location.
 3. The method accordingto claim 1, further comprising: based on the value resulting fromprocessing the request reaching a threshold value, signaling theprocessor, by the accumulator, the value threshold being reached.
 4. Themethod according to claim 1, further comprising: based on reaching aqueue length threshold after queuing the accumulation request for laterprocessing, signaling the processor, by the accumulator, the queuelength threshold being reached.